# drafts

name: draft-nsiangani-authenticatedsecuredlayer
rev: 00
title: ASL Authenticated Secure Layer Protocol
time: 2026-02-11T17:04:47Z
group_uri: /api/v1/group/group/1027/
expires: 2026-08-15T17:04:47Z
states: ["/api/v1/doc/state/1/", "/api/v1/doc/state/150/"]
categories: ["Other AI/agent"]
tags: []
fetched_at: 2026-02-27T22:56:09.014962+00:00
abstract:
This document specifies ASL (Authenticated Secure Layer), a
cryptographic
   protocol for establishing an authenticated and encrypted
communication
   session based on a novel arithmetic primitive called phi-ns-x (also
known
   as ns-mod). ASL provides a complete handshake and data transport
layer
   without reliance on traditional algorithms like RSA, Diffie-Hellman,
AES or
   SHA. Instead, it builds on the Phi-Ns cryptosystem, which defines a 
   structured quadratic relation between primes and achieves security
through
   non-linear modular arithmetic and combinatorial complexity. 

   ASL includes an initial authentication and key exchange mechanism
where a
   server public key (composed only of specific allowed elements such as
a
   prime modulus and state parameters) is used by a client to establish
shared
   secrets. The protocol uses a stateful encryption transform (phi-ns-x)
for
   data, with each message round incorporating a public state value that
   changes unpredictably. Shared secret parameters control this state
evolution,
   preventing man-in-the-middle and injection attacks by ensuring that
only
   parties with knowledge of the secrets can produce valid ciphertexts.

full_text:
Network Working Group                                P.J. Nsiangani, Ed.
Internet-Draft                                         ACE Working Group
Intended status: Standards Track                         9 February 2026
Expires: 13 August 2026


                ASL Authenticated Secure Layer Protocol
              draft-nsiangani-authenticatedsecuredlayer-00

Abstract

   Complete ASL draft embedded below as imported text.  Replace this
   abstract with the structured version later if desired.

Status of This Memo

   This Internet-Draft is submitted in full conformance with the
   provisions of BCP 78 and BCP 79.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF).  Note that other groups may also distribute
   working documents as Internet-Drafts.  The list of current Internet-
   Drafts is at https://datatracker.ietf.org/drafts/current/.

   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet-Drafts as reference
   material or to cite them other than as "work in progress."

   This Internet-Draft will expire on 13 August 2026.

Copyright Notice

   Copyright (c) 2026 IETF Trust and the persons identified as the
   document authors.  All rights reserved.

   This document is subject to BCP 78 and the IETF Trust's Legal
   Provisions Relating to IETF Documents (https://trustee.ietf.org/
   license-info) in effect on the date of publication of this document.
   Please review these documents carefully, as they describe your rights
   and restrictions with respect to this document.  Code Components
   extracted from this document must include Revised BSD License text as
   described in Section 4.e of the Trust Legal Provisions and are
   provided without warranty as described in the Revised BSD License.







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Table of Contents

   1.  Imported Internet-Draft Content . . . . . . . . . . . . . . .   2
   2.  References  . . . . . . . . . . . . . . . . . . . . . . . . .  63
   Author's Address  . . . . . . . . . . . . . . . . . . . . . . . .  63

1.  Imported Internet-Draft Content

ACE Working Group                               Author: P.J. Nsiangani Ed.
Internet-Draft                draft-nsiangani-authenticatedsecuredlayer-00
Intended status: Proposed Standard
Expires: September 2026                                   09 February 2026

                  ASL Authenticated Secure Layer Protocol




Status of this Memo

   This Internet-Draft is submitted in full conformance with the
   provisions of BCP 78 and BCP 79.

   This document is subject to the rights, licenses and restrictions
   contained in BCP 78 and BCP 79, and except as set forth therein,
   the authors retain all their rights.

   This document is provided under the terms of the IETF Trust's Legal
   Provisions Relating to IETF Documents (https://trustee.ietf.org/
   license-info) in effect on the date of publication of this document.
   Please review these documents carefully.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF).  Note that other groups may also distribute
   working documents as Internet-Drafts.  The list of current
   Internet-Drafts is at https://datatracker.ietf.org/drafts/current/.

   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet-Drafts as reference
   material or to cite them other than as "work in progress".

Copyright Notice

   Copyright (c) 2026 IETF Trust and the persons identified as the
   document authors.  All rights reserved.

   This document is subject to BCP 78 and the IETF Trust's Legal Provisions Relating to IETF  Documents (https://trustee.ietf.org/license-info) in effect on the date of publication of this document. Please review these documents carefully, as they describe your rights and restrictions with respect to this document. Code Components extracted from this document must include Simplified BSD License text as described in Section 4.e of the Trust Legal Provisions and are provided without warranty as described in the Simplified BSD License.



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   This Internet-Draft will expire in September 2026.

Abstract

   This document specifies ASL (Authenticated Secure Layer), a
cryptographic
   protocol for establishing an authenticated and encrypted
communication
   session based on a novel arithmetic primitive called phi-ns-x (also
known
   as ns-mod). ASL provides a complete handshake and data transport
layer
   without reliance on traditional algorithms like RSA, Diffie-Hellman,
AES or
   SHA. Instead, it builds on the Phi-Ns cryptosystem, which defines a
   structured quadratic relation between primes and achieves security
through
   non-linear modular arithmetic and combinatorial complexity.

   ASL includes an initial authentication and key exchange mechanism
where a
   server public key (composed only of specific allowed elements such as
a
   prime modulus and state parameters) is used by a client to establish
shared
   secrets. The protocol uses a stateful encryption transform (phi-ns-x)
for
   data, with each message round incorporating a public state value that
   changes unpredictably. Shared secret parameters control this state
evolution,
   preventing man-in-the-middle and injection attacks by ensuring that
only
   parties with knowledge of the secrets can produce valid ciphertexts.


Table of Contents

   1. Introduction ................................................... 3
   2. Terminology .................................................... 5
   3. The Phi-ns-x Cryptographic Primitive ........................... 7
      3.1. Fundamental Equations and Structure ....................... 7
      3.2. Sequential Encryption Recurrence (ns-mod) ................. 9
      3.3. Optional Layer Injection and Masking ......................
11
      3.4. Cryptographic Properties of Phi-ns-x ......................
12
   4. ASL Protocol Specification .....................................
14



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      4.1. Roles and Key Components ..................................
14
      4.2. Session Establishment (Handshake) .........................
15
      4.3. Round-Based Encrypted Messages ............................
18
      4.4. Message Formats and Encoding ..............................
20
   5. Example Use Cases ..............................................
22
      5.1. HTTPS Without TLS (Web Security) ..........................
22
      5.2. Blockchain Transactions ...................................
23
      5.3. Implicit Client Identity Authentication ...................
24
      5.4. Secure Messaging ..........................................
25
      5.5. Encrypted Tunneling (VPN) .................................
26
   6.Security Considerations ........................................ 27
      6.1. Man-in-the-Middle (MITM) ..................................
27
      6.2. Collision Resistance ......................................
28
      6.3. Secret Leakage and Oracle-Free Design .....................
28
      6.4. Post-Quantum Resilience ...................................
29
      6.5. Unpredictability of States ................................
30
   7. IANA Considerations ............................................
31
   8. References .....................................................
32
      8.1. Normative References ......................................
32
      8.2. Informative References ....................................
33

1. Introduction

   The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT",
   "SHOULD", "SHOULD NOT", "RECOMMENDED", "NOT RECOMMENDED", "MAY",
   and "OPTIONAL" in this document are to be interpreted as described in
   RFC 2119 [RFC2119] and RFC 8174 [RFC8174] when, and only when, they
   appear in all capitals, as shown here.




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   ASL (Authenticated Secure Layer) is a protocol that establishes an
   authenticated and encrypted channel between two parties (typically a
   client and a server) using only modular arithmetic operations for
   cryptographic security. ASL is designed as an alternative to TLS and
   similar protocols, particularly in contexts where post-quantum
security,
   small key sizes, or avoidance of legacy primitives (RSA, ECC, AES,
SHA)
   are desired. The protocol leverages the properties of the Phi-Ns
   cryptosystem [PhiNs] and its sequential encryption primitive (phi-
ns-x) to
   achieve confidentiality and integrity in a novel way.

   The foundation of ASL's security is the Phi-Ns hardness assumption.
Phi-Ns
   is an asymmetric cryptographic primitive based on the structured
   decomposition of a quadratic difference between two primes. In Phi-
Ns,
   given two large primes p and q, the relation q^2 - p^2 = T defines a
   composite gap T (Eq0). This gap T is factored in a controlled way
into a
   triple of secret values (a, b, R), where a and b are small exponents
of
   2 and 3 respectively, and R is a remaining large cofactor. The tuple
   (a, b, R) is then encoded in a randomized manner. An attacker who
only
   knows one of the primes (the public key) and a secure commitment to
the
   encoded structure cannot feasibly recover the other prime or the
tuple,
   because there is no oracle or structural hint to guide the search.
The
   problem reduces to blindly guessing the correct prime and internal
factors
   among an astronomically large set of possibilities. The Phi-Ns
mechanism
   does not rely on integer factorization or discrete logarithm; instead
   its security arises from the ambiguity and combinatorial explosion of
   reconstructing the hidden structure from minimal public information.

   ASL uses this primitive in a session context. The protocol is
composed
   of two phases: an **authentication and key exchange phase**
(handshake),
   followed by a **secure data transport phase**. During the handshake,
the
   server proves its identity and the two parties establish shared
secret



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   parameters that will drive the encryption of messages. The handshake
   avoids the need for any external certificate signatures or third-
party
   trust anchors beyond a one-time publication of the server's public
key
   (and an associated commitment value). The client authenticates the
server
   by verifying a static masked value derived from q^2 (the servers
hidden
   prime square), ensuring the server possesses the corresponding
secret.
   Once authenticated, the client and server use phi-ns-x to exchange or
   derive symmetric secrets such as trajectory steps for the encryption
   state. In the data transport phase, every message is encrypted using
a
   **sequential modular transform**: each encryption round employs the
   phi-ns-x recurrence with evolving public state values. Only the
legitimate
   parties (having the shared secrets from the handshake) can maintain
and
   predict the correct state sequence, making it infeasible for an
outsider
   to inject valid messages or decipher the traffic.

   Unlike TLS which depends on PKI certificates, symmetric ciphers
(AES), and
   hash-based HMAC for integrity, ASL achieves all these objectives
through
   pure arithmetic. Authenticity is ensured by the possession of secret
   primes/factors and the correct derivation of state; confidentiality
is
   provided by the non-linear encryption transform; and integrity is
   implicitly guaranteed since any modification to a ciphertext or state
by
   an adversary results in a decryption failure (the decrypted data will
not
   match the expected structure, and the receiver can detect this). The
   protocol is suitable for environments where conventional cryptography
is
   hard to deploy: e.g., IoT devices with constrained resources (small
keys
   and low processing overhead), or specialized scenarios like
blockchain
   transactions where on-chain data must be encrypted yet verifiable.

   The remainder of this document is organized as follows. Section 2
   defines terminology used in the specification. Section 3 describes
the



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   phi-ns-x cryptographic primitive in detail, including its
mathematical
   definitions (Eq0 through Eq6) and cryptographic properties (non-
linearity,
   entropy, direct non-factorizability). Section 4 specifies the ASL
protocol:
   the roles of client and server, the format of public keys and
messages,
   the handshake exchange (including the use of a q^2_masked value for
static
   signature), the concept of a rotating state in encrypted messages,
and the
   handling of shared secrets. In Section 5, we present several use
cases
   to illustrate how ASL can be applied: secure web communication
without TLS,
   blockchain with encrypted payloads, implicit client identity
verification
   via key usage, end-to-end secure messaging, and encrypted tunneling.
   Section 6 discusses security considerations, including resistance to
MITM,
   collision and leakage analysis, post-quantum aspects, and the
unpredictability
   of protocol state. Finally, Section 7 notes IANA considerations
(none) and
   Section 8 provides references to relevant literature and standards.

   By introducing ASL as a formal Internet-Draft, our goal is to enable
   further scrutiny and implementation of this protocol. The design
emphasizes
   a self-contained approach to security where all critical components
are
   derived from a single hard problem (the phi-ns structured gap) and
its
   associated modular arithmetic transform. This approach can complement
   ongoing efforts in post-quantum cryptography by providing a session-
layer
   construction that is not only resistant to known quantum attacks but
also
   efficient and compact.

2. Terminology

   This section defines terms and notation used throughout this
document.
   All definitions are normative for the ASL protocol and the underlying
   phi-ns cryptographic primitive.




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   **Prime**: A positive integer greater than 1 whose only divisors are
1 and itself.

   **p**: A secret prime in the Phi-Ns cryptosystem. In the PK-q mode of
operation
   (public key = q), p is the private key. In the PK-p mode (public key
= p),
   p is public while q is kept private. p participates in the core
equation q^2 - p^2 = T.

   **q**: A prime associated with p in the Phi-Ns system. In PK-q mode,
q is the public
   key; in PK-p mode, q is private. q is chosen such that q^2 - p^2 = T
for some composite T.

   **T**: The quadratic gap defined by T = q^2 - p^2 (Eq0). T is a
strictly positive composite
   integer resulting from the difference of the two prime squares. By
construction, T is
   divisible by 4 (since q^2 - p^2 = (q - p)*(q + p), and q and p are
either both odd or both even;
   for primes >2, both are odd, so q-p and q+p are even, making T
divisible by 4).

   **a**: The exponent of 2 in the factorization of T. Specifically, a
is the largest non-negative
   integer such that 2^a divides T, but 2^(a+1) does not. The value `a`
is part of the secret
   tuple and is not revealed publicly.

   **b**: The exponent of 3 in the factorization of T (after removing
the 2^a factor). Specifically,
   b is the largest non-negative integer such that 3^b divides T / 2^a,
but 3^(b+1) does not. This
   value is also secret (part of the tuple) and not revealed. (Note: In
the context of the encryption
   recurrence, the symbol b_i will be used to denote a different concept
"door" values for injection
   which is unrelated to this exponent b. The usage will be clear from
context.)

   **R**: The remaining composite factor of T after extracting 2^a *
3^b. That is, R = T / (2^a * 3^b).
   R is an odd composite (not divisible by 2 or 3) and may consist of
one or several prime factors.
   R is secret.

   **abR**: The structured decomposition of T into the tuple (a, b, R).



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This triple encapsulates all
   secret factors derived from the quadratic gap. Knowledge of abR
(along with whichever prime was
   kept private) is equivalent to knowing the full private key.

   **Serialized abR**: A randomized encoding of the (a, b, R) tuple. In
Phi-Ns key generation, the atoms
   (prime factors) comprising R are first partially factored (e.g.,
small factors extracted), then all
   resulting factors ("atoms") are permuted in a pseudorandom order and
grouped into blocks. The final
   encoded form includes the values a, b, a salt, the number of blocks,
and the concatenated blocks of
   atoms in permuted order. This serialized abR is committed (e.g., via
a hash or a public checksum)
   to serve as a validation of the key without revealing structure.

   **Recursive decomposition**: An optional process in Phi-Ns where the
secret prime p itself is
   recursively defined by a smaller instance of the problem. For
example, p^2 = p2^2 + T2, with its
   own (a2, b2, R2) decomposition. Recursion can be applied multiple
times, each time increasing the
   entropy of the private key without increasing public key size. If
recursion is used, the abR tuple
   implicitly contains nested layers of factors. ASL and phi-ns-x
operate at the top level; recursion
   mainly affects key generation and security strength.

   **PK-q**: A public key exposure mode in which q is the public key and
p (along with abR) is kept
   private. Attackers see q and a commitment to the serialized abR. They
must attempt to guess the
   correct p (and structure) that yields that commitment. This is the
"standard" mode analogous to
   most public-key systems where the public key is used for encryption.

   **PK-p**: A public key exposure mode where p is public and q (with
abR) is private. Attackers see p
   and the commitment, but have no direct information about q or T. This
mode offers higher security
   because the attacker does not even know the magnitude or sign of T =
q^2 - p^2; q could be slightly
   larger than p or vastly larger, introducing more uncertainty.
However, since encryption operations
   usually require knowledge of q (as an anchor for computations), PK-p
mode in ASL is supported by
   providing alternate public parameters (such as q^2_masked and related



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values) that allow encryption
   without revealing q.

   **M**: A large prime modulus under which phi-ns-x sequential
encryption operations are performed.
   All arithmetic for message encryption in ASL takes place in the
finite field Z_M. M is chosen such
   that it is large enough to accommodate message space and security
requirements (often on the order
   of 2^N for some N, or a safe prime). M is public and is one of the
parameters included in the
   server's public key profile.

   **g**: A generator of the multiplicative group Z_M^*. The value g is
used to advance public states
   in a cyclic manner. Typically, g is a primitive root modulo M. g is a
public parameter.

   **u_k**: The public round state for round k. This is an element of
Z_M^* (i.e., 1 <= u_k < M). In each
   round of encryption, a fresh u_k is used. The sequence {u_k} is
deterministic for those who know the
   secret trajectory parameters (like step sizes), but appears
pseudorandom to others. The initial state
   u_0 (for round 1) is provided by the server during the handshake.
Subsequent u_k are either transmitted
   with each message or recomputed by each side using the shared secret
state update rule.

   **a_k**: The public phase for round k, derived from u_k. It is
computed as
   a_k = (u_k + inv(u_k)) * inv2 (mod M) (Eq2), where inv(x) denotes the
modular inverse of x in Z_M,
   and inv2 is the modular inverse of 2 mod M. The value a_k lies in Z_M
and is used in the encryption
   transform. The pair (u_k, a_k) can be viewed as public state
parameters for the encryption of that
   round's message. Note that given u_k, anyone can compute a_k easily
(so it may not always need to be
   explicitly transmitted if u_k is known).

   **step_u**: A secret trajectory parameter controlling how the public
state u evolves each round. For
   example, a simple state update rule is u_{k+1} = u_k * g^{step_u}
(mod M). The integer step_u (which
   may itself be randomly chosen or derived from the private key
structure) is known to both client and
   server after the handshake but is never revealed publicly. Without



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knowledge of step_u, an attacker
   cannot predict future states from past ones.

   **door values (b_i)**: Optional per-layer additive constants injected
into the encryption recurrence
   to enforce a particular secret structure or order. In some profiles,
each layer i of the multi-layer
   encryption uses an added term b_i (mod M) derived from the secret
decomposition (for example, from
   factors of R or their permutation). These values are kept secret
(shared only between sender and
   receiver) and if used, they modify the encryption formula. The term
"door" implies that these values
   act as additional one-way doors; without knowing the correct b_i
sequence, an attacker cannot properly
   encrypt or decrypt. (This use of b_i is distinct from the exponent b
in the abR structure.)

   **stars**: A conceptual term referring to internal atomic factors or
values derived from the secret
   structure which may be used to randomize encryption. For instance,
the prime factors of R (after some
   processing) can be thought of as "stars". Star values might be used
to derive door values or to vary
   certain steps of the encryption (such as choosing different q-related
multipliers per layer). Some
   deployments publish *star projections*, which are transformations of
these secret star values (e.g.,
   g^(star) mod M), in a shuffled order. The actual usage and ordering
of stars in the encryption process
   remains secret (determined by parameters like star index and
rotation), so an adversary may see a set
   of possible contributions but cannot know which were applied when.

   **rotation**: In the context of ASL, rotation refers to a secret
offset or scheduling parameter that
   determines how internal factors (stars or others) are rotated or
selected across encryption layers or
   rounds. For example, a rotation parameter `rot0` might set an initial
starting index into a list of star
   factors, and with each layer or message, the index shifts according
to another secret step (step_star).
   Rotation ensures that even if multiple internal factors are known in
aggregate (e.g., via published star
   projections), the sequence in which they influence the encryption is
unpredictable to an attacker. The
   rotation parameters are shared secrets between client and server.




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   **round**: A single encryption and transmission event in an ASL
session. Each round k uses a unique
   u_k (and corresponding a_k) and results in a ciphertext (or set of
ciphertext blocks) that constitutes
   one message. A session will have an ongoing round counter. Rounds may
increment for every message
   sent in one direction, or they may alternate between parties
(depending on the implementation, e.g.,
   client-to-server messages could use odd rounds and server-to-client
even rounds). In any case, both
   sides keep track of round numbers and associated state to remain in
sync.

   **q^2_masked**: A public value included in the server's key profile
which serves as an *authenticated
   commitment to the value of q^2*, without revealing q. This value acts
like a static signature or
   identity binding for the server's hidden prime. For example,
q^2_masked might be computed as a certain
   residue or combination: such as q^2 modulo some large public number,
or q^2 with certain parts obscured
   by secret factors. The design of q^2_masked is such that an attacker
cannot derive q from it, but a
   legitimate client (who has obtained the correct public key parameters
out-of-band or from a trusted
   directory) can use it to verify that the server's provided parameters
are consistent with the genuine q.
   In ASL, q^2_masked is used during the handshake to prevent man-in-
the-middle impersonation: it MUST
   match the expected value for the server's identity, or else the
client aborts the connection.

   **ASL session**: A run of the protocol between a client and a server,
from the initial handshake through
   the exchange of data rounds, terminating either by explicit closure
or timeout. An ASL session yields
   a set of shared secrets (trajectory parameters, etc.) and an evolving
state, which are used to encrypt
   and decrypt messages between the parties.

3. The Phi-ns-x Cryptographic Primitive

   Phi-ns-x (also referred to as the ns-mod transform) is the sequential
encryption core used by ASL.
   It operates in a multiplicative group of integers modulo a prime M.
This section describes the
   mathematics of phi-ns-x in detail, including the derivation of its
equations (Eq0 through Eq6) and



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   an explanation of its cryptographic properties.

   At a high level, phi-ns-x can be seen as a keyed non-linear
transformation that takes an element
   (representing a plaintext or intermediate state) and produces another
element (ciphertext or next
   state) such that reversing this transform is infeasible without
knowledge of a secret (the prime
   p or equivalent internal parameters). Each transformation can be
iterated (layered) to amplify
   confusion and hide any algebraic relationship between input and
output. The design is inherently
   modular and does not rely on any polynomial or linear structures that
could be exploited by an
   attacker.

3.1. Fundamental Equations and Structure

   We begin with the fundamental equation defining the relationship
between the two primes p and q,
   and the composite value T:

   **(Eq0)**: q^2 - p^2 = T

   This equation is the cornerstone of the Phi-Ns cryptosystem. T must
be non-zero, positive, and composite.
   Given Eq0, one can express T as a difference of squares: T = (q -
p)*(q + p). However, simply knowing T
   (if it were published) would trivially leak information about p and q
(since one could attempt to factor
   T or solve the equations for p and q). Therefore, T itself is never
revealed. Instead, the structure of
   T is hidden through factorization and encoding.

   The first step in handling T is to factor out small prime powers 2^a
and 3^b:

   **(Eq1)**: T = 2^a * 3^b * R

   where `a` = ord_2(T) (the exponent of 2 in T), and `b` = ord_3(T /
2^a) (the exponent of 3 after removing
   the 2-power). After extracting these, R = T / (2^a * 3^b) is the
remaining factor. By definition of a and b,
   R is not divisible by 2 or 3. R could be prime or composite; if
composite, it may have several prime factors
   which are typically of varying sizes.

   The tuple (a, b, R) is the raw "structured secret" derived from T.



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All parts of this tuple are kept secret
   in Phi-Ns. In practice, the values a and b are relatively small (they
count small factors), while R carries
   most of the cryptographic weight (bits of entropy).

   Next, Phi-Ns employs a **randomized serialization** of (a, b, R) to
ensure that no structure of R or the
   order of factors leaks. This process can be summarized as:

   - Partially factor R (not revealing full factorization if R is large,
but extracting small factors up
     to a certain bound). Each factor (whether prime or an unfactored
cofactor) is considered an "atom".
   - Generate a random seed and salt for this key.
   - Permute the list of atoms using a pseudorandom permutation derived
from the seed.
   - Partition the permuted atoms into a certain number of blocks (the
number of blocks may itself be derived
     from a pseudorandom function of the seed or salt).
   - Construct the serialized encoding as: [a, b, salt, block_count,
(block1 || block2 || ... || blockN)],
     where || denotes concatenation of the binary representation of each
block. Each block might be delimited
     by length or another indicator so that the parsing is unambiguous.

   The serialization is deterministic given the seed, but appears random
to anyone without the seed. Even if
   an attacker could factor R (which might be extremely large or
structured to resist factoring), the permutation
   and block partitioning mean the attacker cannot tell which factors
correspond to which positions in the original
   T structure. In essence, many different tuples (a, b, R) could result
in the same multiset of factors, and the
   serialization further obscures the arrangement. The total number of
possible encodings of the same T (due to
   permutations and partitions) is combinatorially large, adding another
layer of difficulty to any brute force attack.

   In implementations, the final serialized abR is typically committed
via a cryptographic hash (e.g., a SHA-256
   digest) or another one-way function, and that commitment is published
as part of the public key. In ASL, we
   avoid external hash functions as a core dependency (per the design
goal of pure arithmetic security), so an
   alternative approach is used: the q^2_masked value and possibly other
arithmetic check values serve to commit
   to this structure. Still, conceptually, one can imagine that the
client is assured of the integrity of abR



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   through these mechanisms, meaning the server cannot change its
internal structure without detection once the
   public key is set.

   To summarize, the public key information made available (depending on
mode) is:
   - Either p or q (one prime, not both).
   - A commitment to the serialized abR (through q^2_masked or similar
means).
   - Possibly, parameters like M, g, and some initial state or other
auxiliary public values.
   The attacker is missing one prime and the entire abR tuple. The
security assumption is that finding the missing
   values is infeasible.

3.2. Sequential Encryption Recurrence (ns-mod)

   Phi-ns-x defines how a message is encrypted in a sequential (layered)
fashion using the public state parameters
   (u_k, a_k) and the secret prime (with its abR structure). All
operations take place in the finite field Z_M.

   Let us consider a single round k of encryption (for a given session
state). We have:
   - Public inputs: u_k in Z_M^*, and a_k computed as a_k = (u_k +
inv(u_k)) * inv2 (mod M). (We restate Eq2 here.)
   - A plaintext message represented as an integer m in the range [0,
M-1] (if m is larger or is a byte sequence,
     it can be parsed into one or multiple integers mod M or handled
with padding; for now assume m fits in Z_M).
   - The encryption will be applied in L layers (iterations) to increase
non-linearity. L may be a fixed parameter
     or derived from the private key (for instance, it could correlate
with the depth of factorization or be a
     security parameter indicating how many times to apply the
recurrence).

   We denote p_0 = m (the plaintext as the initial state for
encryption). Then for each layer i from 1 to L, we compute:

   - **(Eq3)**: y_i = (p_{i-1} + a_k) mod M.

   This is a simple shift of the previous state by the phase a_k. Note
that if p_{i-1} = -a_k (mod M), then y_i = 0;
   in general encryption we must ensure y_i != 0 (mod M) because we will
require a modular inverse of y_i in the next step.
   If y_i does turn out to be 0, this indicates a degenerate case for
this round (a collision between -a_k and the state);



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   the specification dictates that in such a case the encryptor SHOULD
abandon this round and either increment to the next
   round (i.e., update u and a to new values) or apply an agreed-upon
remediation (such as adding a small nonce to m and
   trying again). This situation is extremely unlikely for random a_k
and p_{i-1} unless an attacker deliberately crafts
   m = -a_k, but we mention it for completeness.

   Given y_i (which is non-zero mod M), define inv(y_i) as the
multiplicative inverse of y_i modulo M. The core
   encryption transform for layer i is then:

   - **(Eq4)**:
     ```
     p_i = ( (y_i^2 - y_i) + (p_{i-1} * a_k) ) * inv(y_i)  (mod M).
     ```

   Expanding and simplifying Eq4, we can see the transformation more
clearly:

     p_i = [y_i^2 - y_i + p_{i-1} * a_k] * y_i^{-1} (mod M).

   This can be interpreted as a quadratic recurrence. If we substitute
y_i = p_{i-1} + a_k, then:
   y_i^2 - y_i = (p_{i-1} + a_k)^2 - (p_{i-1} + a_k) = p_{i-1}^2 +
2*p_{i-1}*a_k + a_k^2 - p_{i-1} - a_k.
   So the numerator in Eq4 becomes: p_{i-1}^2 + 2*p_{i-1}*a_k + a_k^2 -
p_{i-1} - a_k + p_{i-1} * a_k.
   Simplifying: = p_{i-1}^2 + (2*a_k + a_k)*p_{i-1} + a_k^2 - p_{i-1} -
a_k.
   = p_{i-1}^2 + 3*a_k * p_{i-1} + a_k^2 - p_{i-1} - a_k.

   However, this expansion is not particularly illuminating for
cryptanalysis; what matters is that the update
   from p_{i-1} to p_i is highly non-linear (it involves a quadratic
term p_{i-1}^2 and cross-term p_{i-1} * a_k,
   all under modular arithmetic). The division by y_i (which itself
depends on p_{i-1}) further entwines the terms.

   After L layers, we obtain p_L, which we denote as X (the result after
encryption layers). If L=1, we have a single
   application of the transform; typically L might be on the order of 2
to 5 for practical parameters, but it can be
   larger if needed. The final value X = p_L is the output of the
sequential encryption for this round.

   It's important to note that the encryptor, to compute this, only
needs to know:



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   - a_k (publicly computable from u_k).
   - The value u_k (public, but not directly used in formula except to
get a_k).
   - The plaintext m (as p_0).
   - The number of layers L and if any other per-layer parameters are
fixed by protocol.

   The encryptor does **not** need to know p, q, or any of the secret
decomposition (those are only needed for key
   generation and possibly for the receiver to verify or derive some
things). This is intentional: it means anyone
   with the public parameters and current state can encrypt a message.
In ASL, however, not just anyone should be
   allowed to encrypt data that the server will accept; we mitigate
unauthorized encryption through the use of
   shared secret parameters (like door values or required structure in
the ciphertext) described later. But
   mathematically, the transform itself is public.

   On the decryptor side, if one has the secret prime p (and knowledge
of a_k and X), recovering p_{L-1} (and eventually
   m = p_0) is possible by inverting the recurrence using p and the
internal key structure. The inversion is non-trivial
   for an outsider but straightforward for someone with the correct key:
essentially, the decryptor will run a process
   to find which candidate p_{L-1} could produce X under the transform,
guided by the secret knowledge (like knowing q
   or certain invariants). This process is beyond the scope of this
section, but suffice to say it exists and is efficient
   for the legitimate party. For an attacker, trying to invert without p
is akin to solving a random quadratic equation
   in a large finite field with no contextual cluesa problem believed to
be infeasible for cryptographically chosen sizes.

   For completeness, we note that some algebraic rearrangement can
express Eq4 in alternative forms. For example, one can
   multiply out the division to get:
p_i * y_i  y_i^2 - y_i + p_{i-1} * a_k   (mod M),
or:
p_i * (p_{i-1} + a_k)  (p_{i-1} + a_k)^2 - (p_{i-1} + a_k) + p_{i-1} *
a_k   (mod M).
Simplifying that:
p_i * (p_{i-1} + a_k)  p_{i-1}^2 + a_k^2 + 2*p_{i-1}*a_k - p_{i-1} - a_k
+ p_{i-1}*a_k   (mod M),
which reduces to:
p_i * (p_{i-1} + a_k)  p_{i-1}^2 + 3*p_{i-1}*a_k + a_k^2 - p_{i-1} - a_k
(mod M).
There is no obvious further simplification or factorization here; this



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is a purposely irreducible-looking relation
mod M, which is good for security.

In summary, the sequential encryption recurrence (Eq3 and Eq4) takes an
input p_{i-1} and produces p_i under control
of "keyed" value a_k. For a full encryption round of a message, we
iterate this L times (using the same a_k each layer
in a given round k).

**After L layers**, the final output X = p_L is either used directly as
the ciphertext of round k or undergoes a final
post-processing to bind it to the identity parameters (see below). For
many scenarios, X itself is transmitted as the
ciphertext.

If we make an analogy, one can think of phi-ns-x as a sort of "stream
cipher" or "state cipher" where a public state
(u_k, a_k) acts somewhat like an IV (initialization vector) and the
secret prime p and its hidden structure act
like the key that allows decryption. However, unlike a conventional
stream cipher IV, here the IV (state) changes
in a correlated way each round under secret control (step parameters),
and unlike a normal block cipher, the
transformation is heavily non-linear with a trapdoor.

3.3. Optional Layer Injection and Masking

The base transform described by Eq3 and Eq4 is already non-linear.
However, to further thwart any attempt at
cryptanalysis, some deployments of phi-ns-x introduce additional secret
inputs per layer. We refer to this as
"door injection" because it introduces extra secret "doors" that one
must know to successfully navigate the encryption
path.

Concretely, if we have secret door values b_i for layer i (with 1 <= i
<= L), the encryption recurrence can be
modified to:

- **(Eq5)**:
  ```
  p_i = ( (y_i^2 - y_i + p_{i-1} * a_k) * inv(y_i) + b_i )  (mod M).
  ```

Compare Eq5 with Eq4: we simply add b_i (mod M) to the result of the
core transform. The value b_i is chosen in some
deterministic but secret way, possibly derived from the abR structure.
For example, if R has multiple prime factors,



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each layer's b_i could be linked to one of those factors or some
function of it. Another approach is to derive all b_i
from a pseudorandom generator seeded with some secret (like t0 in the
blockchain example, which might be a random or
chosen starting value).

If door values are used, the precise method by which they are generated
and applied MUST be agreed upon as part of
the protocol profile. This document specifies the extension point but
not the exact derivation, as it can differ
based on use case:
- In one profile, b_i might all be zero (meaning no injection, which
defaults to Eq4 behavior).
- In another, b_i might cycle through a small set of constants derived
from the secret key.
- In an advanced usage, b_i might be tied to "star factors": e.g., if R
had factors r1, r2, r3..., we might define
  a series where b_1 = f(r_{idx0}), b_2 = f(r_{idx0+1 mod n}), etc.,
effectively rotating through factors with some
  function f (like taking a reduced residue). The `rot0` and `step_star`
parameters (see Terminology) would control
  how the index through star factors advances with each layer or round.

The effect of Eq5 is that even if an attacker somehow deduced the
functional form of Eq4, they would still be missing
an additive piece in each layer. Without knowing b_i, an attackers
attempt to guess p_i from p_{i-1} and a_k would
systematically fail, as they'd always be off by this unknown offset.
Only the true receiver, who can either compute
or brute-force check the correct b_i (knowing the key, there might be
only one value that makes the decrypted plaintext
intelligible or fits the key structure), can properly invert the
equation.

For interoperability, ASL defines that by default no layer injection is
used (b_i = 0 for all i) unless a specific
profile explicitly incorporates it. This keeps the core specification
simpler. However, implementations SHOULD support
at least the possibility of one non-zero door value (e.g., a constant
offset or a simple pattern) to be future-proof
or to integrate with extended profiles that enhance security via this
mechanism.

Additionally, beyond per-layer injection, one may consider a final
masking of the ciphertext. One such method,
inspired by the Phi-Ns identity binding, is to incorporate q (or T) into
the output. For instance:




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- **(Eq6)**:
  ```
  C = (X + q)^2 - T * a_k^2,
  ```
  where X = p_L is the output after L layers. In this construction, the
transmitted ciphertext C is a quadratic
  residue that depends on q and T. The receiver, knowing T (since they
know p and q), can compute (X + q)^2 = C + T * a_k^2,
  and since they know q, they can solve for X (because (X+q)^2 is known,
taking square root yields two options (X+q),
  but one of those will be consistent with the expected sign or range of
X+q, etc.). This effectively uses the knowledge
  of T as an extra key: an attacker who intercepts C cannot recover X
without solving the equation involving T (which
  is secret). If M and the size of q are large enough, even trying to
brute force q from q^2 (mod something) would be
  infeasible, especially since a_k^2 * T is subtracted which acts as a
blinding factor for outsiders.

The Eq6 example is just one possible "ciphertext projection". ASL may
employ simpler approaches for final output as well,
such as transmitting X as-is, or X plus some constant mask. The choice
can be left to profile definitions. The goal of any
such masking is to tie the ciphertext to the static key in a way that
doesn't affect legitimate decryption but makes
cryptanalysis harder.

3.4. Cryptographic Properties of Phi-ns-x

The phi-ns-x transform possesses several properties that underpin the
security of ASL:

- **Non-linearity**: The core equations (Eq3Eq5) are non-linear
(quadratic) in the plaintext/state variable. Unlike
  linear transformations (which could be solved or inverted with linear
algebra or by observing linear relations),
  phi-ns-x produces outputs that are quadratic residues of input
combinations. The presence of terms like p_{i-1}^2
  and p_{i-1} * a_k inside a modulus means the transformation does not
have a simple algebraic inverse without knowing
  secret parameters. Cryptanalytic attacks that rely on linearity, such
as certain forms of differential or linear
  cryptanalysis, find no straightforward linear approximation here.

- **High Entropy Mixing**: Each layer of transformation mixes the
previous state with the phase a_k and (if used) door
  value b_i. Even if the plaintext has a structure or the attacker can
choose plaintexts, the output after one layer is



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  significantly de-correlated from the input. After L layers, the output
is effectively psuedo-random from the attacker's
  perspective (assuming a_k is not under attacker controlwhich it isn't,
since it's derived from a secret-driven state
  u_k). The introduction of multiple layers and optional per-layer
secrets ensures that even a single round of encryption
  in ASL provides a high-entropy ciphertext.

- **Combinatorial State Space**: If an attacker attempts to guess the
secret (whether it's p or the abR structure or the
  door values), they face a combinatorial explosion. For example,
consider the problem of forging a valid ciphertext
  without knowing the secrets. The attacker would have to guess the
correct p_{0} (plaintext) that, under the unknown
  transform, yields a meaningful result, guess the unknown b_i for each
layer if any, and produce an output that the
  receiver will accept. Because any wrong guess is effectively
undetectable to the attacker (there is no oracle giving
  feedback if a guess is "close" or partially correct), the attacker is
left with an all-or-nothing brute force. This
  is akin to the situation with the underlying Phi-Ns key: an attacker
guessing p gets no feedback that helps refine
  the next guess.

- **No Direct Factorizability**: The encryption transform does not lend
itself to being "factored" or decomposed into
  independent sub-problems. In some ciphers, an attacker can break the
problem into parts (for instance, attacking
  rounds one by one, or separating a key into two halves via meet-in-
the-middle). In phi-ns-x, the structure is such that
  any partial knowledge (say an attacker hypothetically guesses part of
the state or one layer's output) does not linearly
  reduce the complexity of determining the rest. The presence of non-
linearity and modulus coupling means that solving
  for any unknown inherently requires solving the whole set of
equations. This holistic complexity is by design; it's
  analogous to the way that in Phi-Ns key recovery, one must guess
everything (p and full abR) to verify a solution, with
  no way to systematically solve piecewise.

- **Entropy Amplification via Recursion**: If recursive decomposition is
used in key generation, the effect on security
  is an exponential increase in search space for an attacker. For the
legitimate participants, recursion is mostly
  transparent (they just have a larger key to work with, but
encryption/decryption operations remain the same top-level
  ones). For an attacker, however, recursion means that even if they
tried to brute force something like p or q, that



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  would not be sufficientthey would face another hidden layer of
structure inside p itself, and so on. The security of
  ASL with recursively structured keys thus compounds with each layer of
recursion.

- **Post-Quantum Resistance**: There are no known quantum algorithms
that can solve the Phi-Ns inversion problem in
  polynomial time. Shor's algorithm, which breaks RSA and ECC, is
ineffective here because it requires working with a
  known composite or performing an order-finding in a group derived from
known primesPhi-Ns reveals no such composite
  and no such group structure to the attacker. The best one could do
with a quantum computer is apply Grover's algorithm
  to brute force search for the secret, which gives a quadratic speedup.
However, because the search space (the set of
  possible p and abR combinations consistent with the public info) grows
combinatorially with key size (especially due
  to randomization and recursion), even Grover's quadratic speedup is
insufficient to threaten realistic parameter sizes.
  For example, if the classical brute force space is 2^256, Grover might
reduce that to 2^128 steps, which is still
  computationally infeasible.

- **No Partial Disclosure or Oracle**: The design of ASL and phi-ns-x
ensures that an attacker cannot gain partial
  information that leads to a snowball effect. The protocol never
discloses T, the exact abR ordering, or any one-way
  function output that could be inverted. There is also no encryption or
decryption oracle accessible to attackers:
  during handshake, if an attacker tries to impersonate or send
incorrect values, the protocol will simply fail (the
  client or server will abort), without giving the attacker clues as to
which part was wrong. During data transmission,
  an attacker cannot query the decryptor in any useful wayif they send a
bogus ciphertext, the receiver will either
  drop it (if authenticity checks fail) or decrypt to nonsense. This "no
oracle" principle is inherited from Phi-Ns and
  is preserved in ASL. It is crucial for preventing adaptive attacks; an
adversary cannot, for instance, tweak a ciphertext
  slightly and ask "did that decrypt correctly?" as a means to learn
about the key.

In summary, the phi-ns-x primitive provides a robust cryptographic core
for ASL. It ensures that without the private
key material (p and abR, plus any shared trajectory secrets), an
attacker faces a cryptographically hard problem to
either decrypt messages or forge new ones. The combination of structural
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(from phi-ns-x's state evolution) makes ASL a moving target,
cryptanalytically speaking.

4. ASL Protocol Specification

In this section we describe the ASL protocol in detail. We start with
the roles and the components that each party
holds, then delineate the handshake (session establishment phase),
followed by the ongoing message encryption process,
and the exact format and encoding of messages.

4.1. Roles and Key Components

ASL involves two primary roles: the **client** and the **server**.
Typically, the server is the party that has a long-term
key pair (the Phi-Ns key) and provides a service (like a website, an API
endpoint, or a blockchain node), and the client
is the party initiating the secure session, who may or may not have its
own long-term key. This asymmetry mirrors usage
patterns of TLS (servers present certificates, clients initiate
handshake). However, ASL can be used in a peer-to-peer
context as well, where both parties have keys and either can act as
"server" for initiating a connection.

**Server Key Materials**: The server possesses:
- A key pair (p, q) along with the secret decomposition (a, b, R).
Depending on the chosen mode:
  - In PK-q mode: q is the public key, and p with abR is private.
  - In PK-p mode: p is public, and q with abR is private.
- A set of public parameters associated with the key:
  - The prime modulus M used for phi-ns-x operations.
  - The generator g for Z_M^*.
  - Possibly one or more recommended profiles or parameter sets (like
the typical size or number of layers L).
- A commitment to the secret structure:
  - Either a cryptographic hash of the serialized abR, or in our design,
a numeric value that serves the same purpose (q^2_masked).
- Optionally, a set of *auxiliary public values* that assist in
encryption without revealing secrets:
  - For PK-p mode: values derived from q or R that allow the client to
perform encryption. For example, **q_anchor** (an "anchor" value related
to q, such as q mod M or q mod some smaller base) and **star
projections** (a list of g^f mod M for some secret factors f). These
were illustrated in the blockchain use-case as `q_anchor` and a list of
star projections. In a general protocol context, these might or might
not be transmitted, but if used, they are included in the server's
initial handshake message.
- Policy information: e.g., how often the key is refreshed (the Phi-Ns



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key can be re-randomized via abR re-serialization without changing p or
q), supported modes (PK-p or PK-q or both), and supported use of
advanced features like door injection.

The server's **public key profile** can thus be summarized as a
structure containing: { mode, p_or_q, M, g, [q^2_masked], [q_anchor],
[star_projections], [others...] }. This profile would be published or
made available to clients via some secure channel (like DNS, a
blockchain, or an offline distribution), similar to how a TLS server has
a certificate.

**Client Key Materials**: In many scenarios, the client may operate
without a long-term key pair (just as most web clients do not have a
certificate). ASL can accommodate a client with no long-term keys; the
client will simply generate some ephemeral values during the handshake
(like a random session secret to send to server). However, ASL also
allows mutual authentication: a client could have its own Phi-Ns key
pair and authenticate itself to the server in a similar way (this would
be analogous to client certificates in mutual TLS, but achieved via the
same mechanism as the servers authentication). For brevity, we focus on
the typical case of server-authenticated, client-unauthenticated
sessions, but we will note where a client key could come into play.

Regardless of having a long-term key, the client will generate some
**ephemeral secrets** during the handshake:
- A random session secret `K_s` (for example, 256 bits) that will be
used to derive shared parameters (like step_u, step_star, etc.). Instead
of directly using this as a symmetric key as in RSA-based TLS, ASL can
incorporate this secret into the phi-ns-x encryption so that only the
server (with its private key) can recover it.
- Potentially an ephemeral Phi-Ns key pair for forward secrecy: The
client might generate a fresh small Phi-Ns instance (p_e, q_e) just for
this session and provide q_e in the handshake (with a commitment to
ensure authenticity). The server could then combine keys in some manner
or simply switch roles and encrypt something back to the client using
the client's ephemeral public key. This is an advanced option and not
required; forward secrecy can also be achieved by discarding the session
secrets after use.

After the handshake, both client and server will share:
- The trajectory parameters for state updates: e.g., step_u (for u
sequence), step_star (for star rotation), rot0 (initial rotation),
star_idx0 (starting index for star usage), L (number of layers for phi-
ns-x), etc. These are either negotiated or derived from the exchanged
secret.
- The current state values: u_0 (initial state for round1, provided by
server), and consequently a_0 (which client can compute).
- Any symmetric key material needed for additional encryption or



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authentication (if we overlay a MAC or use symmetric encryption for
payload, but ASL's design tries to avoid separate MAC by making
ciphertext self-authenticatingstill, a session might derive a key to,
say, authenticate certain out-of-band signals or to mix into application
protocols).
- Optionally, if client authentication was performed, the server
similarly learns some shared secret that confirms the client's identity.

In summary, the server is the holder of the master key (Phi-Ns key) and
provides public parameters; the client initiates communication and
contributes randomness to create unique session keys. The protocol
ensures that after the handshake, both sides have identical sets of
secret parameters which they will use to synchronize the stateful
encryption.

4.2. Session Establishment (Handshake)

The ASL handshake is the process by which the client and server
establish a secure context. The handshake achieves:
1. Server Authentication: The client gains confidence that it is
communicating with the legitimate server (not an impersonator), via the
server demonstrating knowledge of its secret key (p, q) in a non-
forgeable way (using q^2_masked and successful decryption of a test
secret).
2. Establishment of Shared Secrets: The client and server arrive at
common values (the "session keys" or parameters) that will be used for
encrypting and decrypting application data. Only the two parties learn
these values.
3. (Optional) Client Authentication: If required, the server can also
authenticate the client, either via a similar mechanism or by verifying
that the client knows a pre-shared secret. In many applications this
step is omitted, or handled at the application layer (e.g., via a login
after the secure channel is up).

The handshake can be described as a sequence of message exchanges. We
define the messages in an abstract format here:

- **ClientHello**: This is the first message sent by the client to the
server. It may include:
  - A protocol version identifier (to handle future versions or backward
compatibility).
  - A list of cipher suite preferences or capabilities (in ASL's
context, this could indicate whether the client supports certain
profiles, e.g., injection or recursion or PK-p vs PK-q mode
preferences).
  - A random nonce or session identifier (client_random) for the
handshake, to ensure freshness.
  - Optionally, an ephemeral public key of the client if doing mutual



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key exchange (e.g., the client could include q_e (ephemeral) and a
commitment of its own, but this is optional).
  - Optionally, any cookie or anti-DoS token if a prior stateless
negotiation mechanism is used (similar to DTLS or IKE cookies; this is
an operational detail not core to cryptography).
  - The ClientHello in ASL can actually be minimal if the protocol
doesnt need negotiation: in the simplest case, it might just be a packet
indicating "I want to start ASL handshake" plus a random.

- **ServerHello (ServerParams)**: The server responds with its
parameters. This is a critical message containing:
  - The selection of protocol version and profile (if any negotiation;
likely just mirrors what client said if supported).
  - The server's public key information:
    - `mode` (PK-p or PK-q).
    - `p_public` (if mode is PK-p) or `q_public` (if mode is PK-q).
We'll use p_public to denote whichever prime is being revealed as
public.
    - `M` (modulus for encryption).
    - `g` (generator for state updates).
    - `q^2_masked` (the masked commitment to q^2, as defined in
Terminology).
    - If mode is PK-p (so q is secret), then **auxiliary encryption
parameters** must be provided so the client can perform encryption. This
could include:
      - `q_anchor`: a value that serves as a temporary anchor for
encryption calculations. For example, the server might provide q mod M
as q_anchor. Since p is public, q mod M alone might not give away q if M
is much smaller than q or if multiple possible q yield the same mod M
(which can happen if q > M).
      - A set of `star_projections`: e.g., a shuffled list of g^(factor)
mod M for factors of R or related values. This effectively gives the
client a pool of values that relate to secret components, one of which
might need to be used in encryption layers, but the attacker doesn't
know which or how.
      - The rationale is that with p_public known, the client can
compute T mod M if it had q mod M (because T = q^2 - p^2, so T mod M =
(q mod M)^2 - (p mod M)^2 mod M). If M is large and not a factor of T or
anything, this by itself may not leak too much. Anyway, the specifics of
q_anchor usage would be profile-defined. The point is the server
supplies whatever is necessary for the client to execute phi-ns-x
encryption.
    - If mode is PK-q (so q is public), then providing q itself is
enough for encryption; auxiliary parameters might not be needed, though
the server could still provide some precomputed values or
recommendations (like how many layers to use).
  - The server's initial state:
    - `u_0` (or u_1 depending on indexing) for the first round of



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encryption after the handshake.
    - The server might also supply `a_0` for convenience, but since a_0
can be derived from u_0, this is optional. However, including a_0 allows
the client to double-check calculation.
    - A round number or session ID if needed. For example, round=0 might
be implicit for handshake phase.
  - A server_random nonce and possibly a session ID (to pair with
client_random for key derivation and to avoid replay).
  - Optionally, a signature or proof over all these parameters. In
classical TLS, the server would now send a Certificate and a signed
ServerKeyExchange to prove possession of the private key corresponding
to the certificate's public key. In ASL, we achieve this proof not by an
external signature algorithm but inherently:
      * The inclusion of q^2_masked is one form of static proof (the
client will verify this matches the known value for this server
identity).
      * Additionally, the server will prove its private key knowledge by
being able to decrypt the client's next message correctly. If the server
were fake and didn't have the key, it would fail to respond properly in
the next step, which the client would detect.
      * However, there is a potential for a subtle MITM if we aren't
careful: an attacker could relay messages between client and real server
without knowing the key. To prevent that, some binding of handshake
messages is needed, similar to the Finished message in TLS. We will
address that with a final handshake verification.
  - (If client authentication is desired) a request for client key or
info, akin to a CertificateRequest in TLS. In ASL, that could be an
indication like "please prove who you are by sending an encrypted blob I
can verify," but we'll not elaborate here.

- **ClientKeyExchange (ClientEncryptedSecret)**: After receiving the
server's parameters, the client verifies:
  - The server's `p_public` or `q_public` matches the expected identity
(for example, the user might have the server's p or q fingerprint
configured, or in a web scenario, this would be checked against some
known value or a trust on first use store).
  - The `q^2_masked` matches the known commitment for that server (this
might be published through DNS or a public ledger or given out-of-band).
If this check fails, the client MUST abort the handshake (since it
indicates the server's parameters are not consistent with the legitimate
key).
  - If the check passes, the client proceeds to encrypt a secret for the
server. The client generates the random session secret `K_s`. This can
be, for instance, 32 bytes of cryptographically secure random data. The
client then encrypts `K_s` using the phi-ns-x transform under the
provided state and parameters:
    * The client takes the plaintext m = an encoding of `K_s` (if `K_s`
is 256 bits, interpret it as a number modulo M; if its larger, it might



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be split across multiple messages or handled by chunking).
    * The client uses the round state provided (`u_0, a_0`). Let's call
this round "handshake round" or round 0. It will apply the recurrence
Eq3 and Eq4 (and Eq5 if doors are used) L_h times, where L_h might be
fixed or chosen by servers profile. The output will be X_secret.
    * If PK-p mode is used and the server supplied auxiliary parameters
like star projections, the client will incorporate them as specified.
For example, if a door value b_i is derived from a star factor and the
server gave a list of star projections, the client might not know which
to use. However, the safest approach for handshake is not to require the
client to guess any secret injection. So the server could specify that
for the handshake round, no injection is used (or a fixed simple
pattern) just to convey the secret. Alternatively, the server could
supply an explicit hint: e.g., an index or one of the star projection
values that the client should use for encryption in this handshake. To
keep things straightforward, assume that the handshake uses a profile
with no injection (or the server picks parameters such that b_i=0 for
this round). The primary goal is to get `K_s` across.
    * The client sends **ClientEncryptedSecret** = X_secret (the
ciphertext). Additionally, the client may send computed verification
data:
      - It can compute a hash or MAC of the handshake messages
(ClientHello + ServerHello + X_secret) using `K_s` and include this as
well (similar to how TLS's Finished message is a MAC over all prior
handshake messages using the new keys). This serves as a strong
confirmation that the handshake hasn't been tampered with. This can be
done by, say, taking `HMAC(K_s, handshake_transcript)` or a
straightforward hash since we want to minimize external primitives. If
we avoid hash functions entirely, we could use phi-ns-x itself in a
specialized way to produce a tag (e.g., encrypting a constant with a
known pattern using `K_s` as input could produce a tag). But including a
small SHA-256 here purely as a handshake integrity check is not against
the spirit of ASL, since it's not used for core security, but let's
assume we want to avoid even that. Another way is: The server, when it
decrypts `K_s`, could produce a known response that implicitly confirms
the handshake integrity.
      - To keep the initial specification free of additional primitives,
we will proceed without an explicit Finished MAC, but note that this is
a point where a profile could add one for enhanced security.

- **ServerFinish (Handshake Completion)**: Once the server receives the
ClientEncryptedSecret, it attempts to decrypt it:
  - The server knows p (or q in PK-q mode) and the full secret key, so
it can run the inverse of phi-ns-x on X_secret using the same round
state (u_0, a_0) and (if applicable) the known door values. This yields
a candidate plaintext which should be `K_s`.
  - If the server fails to decrypt properly (e.g., because the attacker
tampered with X_secret or it was encrypted under wrong assumptions), the



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result will be gibberish. The server can detect failure perhaps by
format (if `K_s` has some structure or length). If decryption fails, the
server MUST abort the connection.
  - If decryption succeeds and yields a plausible `K_s`, the server now
has the shared session secret.
  - At this point, both sides have `K_s`. They can derive from it the
internal parameters needed for the ongoing encrypted session. A simple
derivation could be:
    * step_u = H1(K_s, client_random, server_random) mod (some range)
where H1 is a hash or KDF function. But again, to avoid hash, they might
directly treat `K_s` bits as comprising these values. For example, split
`K_s` into two halves: first half as step_u, second half as step_star
(just as bits of an integer). If `K_s` is uniformly random, each half is
random. Or derive rot0 from a subset of bits, etc. The exact method
should be agreed: the point is, `K_s` is now the master secret for the
session from which all needed shared secrets are determined.
    * For instance: interpret `K_s` as a 256-bit number. Break it into:
step_u (e.g., 128-bit), step_star (64-bit), rot0 (16-bit), and an
authentication token (48-bit)  these bit lengths are arbitrary examples.
The values may be reduced modulo appropriate ranges (like if step_u must
be taken mod (M-1) maybe, since it might effectively act as an
exponent).
    * Both client and server do the same derivation, so they have
identical parameters.
  - The server should now send a final confirmation. If we had not
included a Finished MAC from client, the server will include something
to assure the client that it too derived the same keys and that the
handshake wasn't tampered with. The simplest way:
    * The server sends a **ServerFinish** message encrypted with the
newly established session keys. For example, the server can compose a
message "HS_DONE" (handshake done indicator) and encrypt it as the first
ASL data round (round 1).
    * Concretely, server sets round=1, computes new u_1 from u_0 (with
secret step_u), computes a_1, encrypts a known ASCII string or
structured data (like a constant or perhaps the handshake transcript
hash) using phi-ns-x with L layers under a_1 (and door values etc. now
known to both from K_s), producing X_verify, and sends it along with
u_1, a_1.
    * The client, upon receiving this, will decrypt using its copy of
the secrets. If the plaintext matches the expected value (e.g.,
"HS_DONE" or matches a hash of handshake messages), then the client
knows the server was indeed present in the handshake and that the shared
secrets match. At this point the handshake is complete and the channel
is secure and authenticated.
    * If the decryption fails or the value is incorrect, the client must
terminate the session because it indicates something went wrong (keys
mismatch or MITM).
  - If we wanted to avoid even one round trip for finish, the clients



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previous message could have included something that server would echo
back. But it's cleaner and more symmetric to have this final step.

- After these exchanges, the handshake is complete. The protocol
transitions to the data transfer phase.

**Summary of Handshake Flow**:
Client -> Server: ClientHello (random_C, cipher_suites, maybe
client_ephemeral_pub)
Server -> Client: ServerHello (p_public or q_public, M, g, q^2_masked,
u_0, [q_anchor, stars...], random_S)
Client -> Server: ClientEncryptedSecret (X_secret = Enc(m=K_s), maybe
Finished MAC)
Server -> Client: ServerFinish (Encrypted confirmation with round 1
state)
(Messages could be combined or split differently in implementations, but
this is the logical content.)

Throughout this process, no out-of-band trust is needed except that the
client must have an authentic reference for
the server's identity (particularly the expected q^2_masked or the
servers public key itself). This could be obtained
via a one-time secure channel or a trusted directory. If that is not
available (trust on first use scenario), the
protocol still encrypts the traffic, but a MITM could impersonate the
server on the first connection. However, because
q^2_masked would change if they later see the real server's key, a
client could detect it on subsequent connections
(similar to SSH's known host keys). This is a deployment consideration.

For clarity, let's tie back to the earlier blockchain example terms:
- `p_public` in handshake corresponds to that p_public.
- `q_anchor` and `q^2_masked` were given in "genesis" in that example;
here ServerHello carries them.
- The shared off-chain secrets like step_u, step_star, rot0 are
essentially derived from K_s in our handshake.
- The round state (u_k, a_k) in each transaction corresponds to our per-
message state below.

4.3. Round-Based Encrypted Messages

Once the handshake has completed, client and server enter the secure
data phase. The key aspect of ASL's data encryption
is that it is stateful: each message is tied to a round index and the
public state (u_k, a_k) for that round.

The participants now share:
- The secret trajectory parameters: step_u, step_star, etc., derived



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from K_s.
- The starting state: let's call it (u_1, a_1) which should have been
either provided in the ServerFinish or computed
  by each side. Actually, in the handshake above, the server already
sent u_0 for the handshake encryption. After using
  u_0, both sides can compute u_1 = u_0 * g^step_u (mod M). If the
server already did this and sent something with u_1,
  then the client upon verifying ServerFinish knows u_1 and a_1.
  Alternatively, some might prefer that the server doesnt send u_1 but
that both sides compute it independently from
  u_0 and step_u. However, if any message was lost or if one side
increments and the other doesn't, there could be desync.
  Its safer to include u_k in each message to avoid synchronization
issues. This is slightly less efficient (sends some
  extra bytes), but ensures robustness.
- The internal secret factors scheduling (like the sequence of door
values b_i if used, or star usage pattern): these are
  agreed or known from handshake secrets.

Data is sent in **records** or **messages**. Each record is self-
contained in the sense that it carries the information
needed for the receiver to process it (or the receiver can derive that
info if in sync).

For each message the following occurs (assuming a message is one
protocol data unit to be encrypted and sent; if a
higher-level application message is large, it could be fragmented into
multiple ASL messages, each with its own round):
1. The sender (could be client or server) determines the next round
number and computes/obtains the corresponding
   state (u_k, a_k).
   - If the sender was the one who sent the last message (i.e., we might
not alternate; consider an HTTP response with
     multiple packets), then the sender would have already incremented
the state internally after sending. Typically,
     we can define that every message increments the round number by 1,
regardless of who sends it. This means both
     sides advance a shared counter even on their own send. That is
unusual (because normally each side would increment
     their own sequence), but since state is shared, it is more like an
alternating sequence in lockstep.
   - Another approach: maintain two independent sequences, one for each
direction (like TLS has independent sequence
     numbers per direction). We could do that: a client_send_state and
server_send_state each with their own step parameters.
     That would require deriving separate step_u_client and
step_u_server from K_s. This might be more elegant to avoid
     issues of one side not sending for a while and the other wanting to



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send many messages (like a server pushing data
     unidirectionally). Yes, it's probably practical to have separate
state threads so that one side sending many messages
     doesn't starve the sequence for the other.
   - Let's assume we have separate sequences: one keyed for client
transmissions, one for server transmissions. In that
     case, we derive step_u_C and step_u_S from K_s. The initial u_C1
and u_S1 can both start at some common value or be
     derived differently (maybe the server provided u_1 in handshake
which was for first server message, and for client
     initial state maybe the client picks or both could use same u but
different step yields divergence).
   - To avoid complicating, though, we can also stick to a single
sequence used by both: in a request-response scenario
     it naturally alternates. But consider asynchronous messaging, a
single sequence is simpler because it globally orders
     messages. This is akin to a half-duplex or synchronized full-duplex
channel. There is a risk if both try to send at
     the exact same round number simultaneously. That could be resolved
by one of them detecting a collision and adjusting
     (e.g., if two messages come with same round from both sides, the
higher party ID wins and the other can resend with
     increment). This is too complex; separate sequences is easier for
practicality.
   - For specification clarity, let's define:
     * Each side maintains a send sequence (with its own u_k evolving
via step_u known to that side).
     * Actually, step_u is shared, so if both used it they'd produce the
same u sequence which is not what we want if
       they are sending concurrently because they'd collide. If we want
separate, we could have step_u_C = step_u (from K_s)
       and step_u_S = -step_u or another independent value. Perhaps we
could derive two seeds from K_s so that one
       sequence might use g^step_u, the other uses g^step_u2 or maybe
g^(-step_u) to differentiate.
     * For simplicity in this version, we assume a synchronous
communication: client sends a request (round 1), server
       responds (round 2), client next (round 3), etc. This fits
scenarios like request-response or blockchain where
       transactions are sequential by design. If a use case needs
simultaneous asynchronous sending, the protocol can
       be extended to label messages with which sequence they belong to
(C or S) and each with its own state update rule.

2. The sender prepares the plaintext for encryption:
   - This plaintext could be a JSON, a binary blob, etc. It may be
necessary to encode a length or ensure it's smaller
     than M. Usually M would be chosen large enough to send a reasonable



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chunk or a symmetric cipher key. But for streaming,
     if data is larger than M can handle at once, one would chunk it.
ASL can either treat each chunk as an independent
     message or incorporate a chaining at the phi-ns-x level (the latter
is not typical; easier to chunk).
   - If needed, the plaintext is padded to a certain length (maybe to
nearest multiple of some bytes) to avoid leaking
     exact lengths. But that's application specific. Not mandated here,
though adding padding for security (to not reveal
     message length patterns) is recommended.

3. The sender encrypts using phi-ns-x:
   - They know the current rounds a_k. They set p_0 = plaintext
(converted to int mod M by some agreed reversible scheme,
     e.g., big-endian byte interpretation).
   - They run the recurrence for L layers (in normal data, L might be
set to a value possibly different from handshake L).
   - If door values (b_i) are being used for data encryption, the sender
applies them as per Eq5. Both sides know the sequence
     of b_i to use. If b_i sequence is static for all rounds, fine.
Possibly, the b_i sequence could actually depend on the
     round number, e.g., using different segments of the star list for
different rounds. That could be orchestrated by using
     rot0 and step_star: e.g., at round k, start injecting factors from
index = (rot0 + k*step_star) mod N. This is a plausible
     design. It means each round's multi-layer encryption uses a
different subset of Rs factors, making patterns between rounds
     even less correlated.
   - After computing X = p_L, the sender may apply final masking if the
profile calls for it. However, typically for sending
     data, we might choose to not use an additional mask (or if we do,
both sides must know how to remove it). If something like
     Eq6 was used, the receiver would need to know it and apply inverse.
For now, assume the ciphertext is X (mod M).

4. The sender transmits the message which includes:
   - The round identifier (k). If we strictly alternate, k might be
implicit (each side expects the next number). In unpredictable
     async, better to include it explicitly.
   - The public state values for that round: u_k and/or a_k. (One of
these could be derived from the other. Possibly sending only
     u_k is enough since a_k can be recomputed. But sending both could
be a nice integrity check: since a_k = (u_k + inv(u_k))*inv2 mod M,
     the receiver can recompute and verify the given a_k matches given
u_k. This could catch any corruption or tampering in transit.
     An attacker who tries to change u_k would also have to adjust a_k
consistently to avoid detection.)
   - The ciphertext X (or C if masked form).



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   - Any additional metadata, e.g., a flag for last message, or an
application-level content type if needed (like TLS records have type).

5. The receiver on receiving a message:
   - Checks the round number (if out-of-order or repeats, might indicate
replay or loss).
   - Verifies the state consistency: e.g., if a_k is provided, check a_k
?= (u_k + inv(u_k))*inv2 mod M. If it doesn't match, the message
     is invalid and should be dropped.
   - If the protocol expects sequential rounds and the round number is
not the expected next (and not yet seen), it might indicate either
     reordering or an attack. Depending on design, might queue out-of-
order or just close connection due to suspicion.
   - Using the locally stored secrets:
     * The receiver knows step_u (or two if separate directions, then
picks appropriate one).
     * If the receiver has been computing states on its own (for its
sending), it might also compute what u_k should be for the peer's
message if in sync.
       If they are using separate sequences, the receiver can compute
expected u for a given round if it knows how the other side's step
works. Actually, if
       separate, maybe the receiver isn't computing the other side's
states on its own, instead it relies on what's in the message. And then
it updates some
       record of "last seen u from that side".
     * In either case, since u_k is included, the receiver doesn't
strictly need to pre-compute it, but it can verify the progression if it
has memory of
       previous u from that sender. For example, if last server message
had u_x, the client can verify that u_k = u_x * g^{step_u_S}^{(k-x)} mod
M (i.e.,
       check it fits the expected sequence progression). If not, either
messages were lost or altered. If lost, maybe we can still accept if
within a window
       and adjust? But let's not complicate; ideally no loss or the
protocol might drop out-of-sync.
   - The receiver then computes a_k from u_k (if not already given or to
cross-check).
   - It sets up the decryption recurrence. Now, decryption of phi-ns-x
is not simply running the inverse function easily because we don't have
a straightforward
     algebraic inverse. But the receiver has the secret p and possibly
other private key info. They need to solve p_L = X for p_0.
     * One approach: Because the legitimate receiver has full knowledge
of p (the secret prime) and the internal structure, it could simulate
the encryption
       forward to see if it matches X, but that would require guess of



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plaintext. Instead, likely the receiver uses a more direct method. For
phi-ns-x, given
       X = p_L, one can iterate backwards:
         - There is a relation linking p_L and p_{L-1}: from Eq4
rearranged, knowing p_L (and therefore presumably knowing y_L since y_L
= something with p_{L-1}),
           it's tricky to go backward directly without knowing p_{L-1}.
However, in the context of the full key, the secret prime p might allow
computing some
           invariant that helps invert layer by layer. This part is
complex; presumably, the design of phi-ns-x ensures decryptability by an
algorithm that leverages
           p and abR. For example, there might be a way to detect the
correct y_i at each step by using the fact that the output must align
with the prime's identity
           structure (something like p_i must be a quadratic residue
related to p or q in a known way).
         - In any case, since phi-ns-x was introduced as an "encryption"
transform with the statement that decryptor possesses p and trajectory
parameters to decrypt,
           we trust that a decryption algorithm exists. Perhaps the
decryptor tries all possible factors as candidate b_i inverses if door
used, etc. Or maybe since it
           knows q, it can compute what y_i should be: e.g., maybe q
acts like an eigen or fixed point parameter such that at final layer L,
(X + q*a_k)^? yields something
           that reveals p_{L-1}. Hypothetically, if we plug p_i = q
(just thinking), then Eq4 would simplify. Actually, maybe the trick: The
server can test candidates
           for p_{L-1} by solving the quadratic equation given p_L and
known a_k and known b_L if any. That quadratic is:
           `p_L * y_L  y_L^2 - y_L + p_{L-1}*a_k - b_L (if b used)`
which is essentially `p_L * (p_{L-1}+a_k)  (p_{L-1}+a_k)^2 -
(p_{L-1}+a_k) + p_{L-1}*a_k - b_L`.
           This can be rearranged to a polynomial in p_{L-1}. Likely a
quadratic in p_{L-1} as well. The decryptor can solve that quadratic
(which might have 0,1, or 2 solutions mod M).
           One of those solutions should be the correct p_{L-1}. How to
choose the right one? Possibly because only one will continue to yield a
valid plaintext in subsequent backward steps or
           within range of message size. Then it can continue
iteratively. This is plausible as a decryption method.
     * We do not need to detail the algorithm in the specification; it's
enough to state that the receiver uses its private key to invert the
phi-ns-x layers and recover plaintext.
       This may involve brute-forcing a small number of possibilities at
each layer (since the equation to invert is quadratic, at most 2
solutions, maybe more with door values considered,



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       but if star injection made ambiguous, the correct combination
must be found).
     * We guarantee that given the design of abR and injection, the
correct plaintext yields a coherent solution chain that matches the
private key structure, whereas a wrong guess quickly
       leads to a contradiction (like an impossible factor or mismatched
q^2 when check against q).
   - After obtaining p_0 (plaintext as an integer), the receiver
converts it back to the original message format (removing padding if
added).
   - The receiver then uses the plaintext or passes it up to the
application.

6. State update: After processing a message, both sender and receiver
update their expected state for that sequence:
   - If using one global sequence, they both increment the round
counter.
   - If separate, the sender obviously updated its own state when
sending. The receiver upon successfully decrypting might recompute what
the next expected u for that sender would be (though it can also just
wait to receive it in the next message).
   - Additionally, if star rotation is used across rounds, both sides
update the star index (e.g., star_idx = star_idx0 + round * step_star
mod N, or some such formula), so that the next message will use a
different starting point in the star list for door values.

Notably, there is no separate "MAC" in each record as one would see in
TLS (prior to TLS 1.3) because the encryption itself combined with state
verification suffices. If an attacker modifies bits in the ciphertext X,
the probability it decrypts to something valid (like meaningful
plaintext with correct structure) is negligible; even if it did, the
protocol's state check (u_k and a_k sequence) would likely fail or the
content would be nonsense to the application and could be detected with
application-level sanity checks. The inclusion of state values provides
an implicit integrity check: an attacker cannot easily alter X without
also altering a_k and u_k in a consistent way, which would require
breaking the encryption anyway.

In practice, an implementation might still want to include an explicit
checksum or MAC over the ciphertext and perhaps part of the header (u_k,
a_k) using the session symmetric secrets, to guard against any
theoretical weaknesses in the transform as an authenticity mechanism.
However, we leave this as optional, given the design goals.

4.4. Message Formats and Encoding

We now specify the format of the handshake and data messages in a more
concrete manner. These formats are described conceptually;



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actual implementations may serialize fields in network byte order, etc.
We assume a transport where messages can be delineated
(e.g., a TCP stream with length prefixes or a datagram protocol with
message boundaries).

**ClientHello Message:**
struct {
uint8 protocol_version_major;
uint8 protocol_version_minor;
uint16 cipher_suites_supported;   // e.g., a bitmask or ID indicating
profile support
opaque client_random[32];        // 256-bit random nonce
// Optional fields
uint8 have_ephemeral_key;        // 0 or 1
if (have_ephemeral_key == 1) {
uint16 ephemeral_key_length;
opaque client_ephemeral_key[ephemeral_key_length];  // e.g., a
serialized phi-ns public key or other
opaque client_ephemeral_commitment[??];            // maybe a q^2_masked
for the clients ephemeral q, if needed
}
} ClientHello;
This is an example. In many cases, cipher_suites_supported might just be
a single profile ID if ASL is simple (for instance, 0 = base, 1 =
withDoors, 2 = PK-p mode only, etc.). The client_random ensures the
handshake has entropy from client side as well.

**ServerHello (ServerParams) Message:**
struct {
uint8 protocol_version_major;
uint8 protocol_version_minor;
uint16 cipher_suite_chosen;
opaque server_random[32];
uint8 key_mode;     // 0 = PK-q (server public = q), 1 = PK-p (server
public = p)
opaque server_public_key<1..+>; // If key_mode=0, this contains q (big-
endian integer). If key_mode=1, contains p.
opaque q2_masked_value<...>;       // A fixed-length encoding of
q^2_masked. For example, same length as M or other standard (must be
defined by profile).
uint16 modulus_length;
opaque modulus_M[modulus_length];
opaque generator_g[modulus_length];  // same length as modulus to hold
value
opaque initial_u[modulus_length];
// if key_mode = 1 (PK-p), include auxiliary data for q hiding:
if (key_mode == 1) {
opaque q_anchor[modulus_length];   // possibly q mod M or some



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representative
uint16 star_count;
opaque star_projections[star_count][modulus_length]; // array of
star_count elements, each modulus_length bytes
// (star_count could be implicitly determined from length too)
}
} ServerHello;
Note: We treat p, q, M, g, u as big-endian byte arrays with a known
length (modulus_length could be, e.g., 32 bytes for a 256-bit prime M).
The star_projections array is an example of optional data; its content
and use must be specified by the profile (the client needs to know what
to do with them, which might be documented in an application profile
spec rather than the core).

**ClientEncryptedSecret Message:**
struct {
opaque encrypted_secret_value<1..+>;  // The ciphertext X (or C) from
encrypting the clients secret.
// Optionally:
opaque handshake_verify_tag[32];       // e.g., a 256-bit tag for
handshake verification (if implemented).
} ClientEncryptedSecret;
The encrypted_secret_value length should equal the length of modulus M
(as ciphertext is an element of Z_M). If the secret is slightly bigger
than M bits, multiple such structures could be sent or it could be
split, but typically we'll choose M to be big enough.

**ServerFinish Message:**
struct {
opaque next_u[modulus_length];
opaque next_a[modulus_length];
opaque verification_ciphertext[modulus_length];
} ServerFinish;
In this, next_u and next_a are state for round 1 (if not already known;
if we assume handshake round was 0, then round 1 is first data round
possibly carrying this verification). The verification_ciphertext is the
encryption of a known value (like all zeroes or some hash of transcript)
using round 1's state. The client will decrypt it and check. If separate
sequences are used, then perhaps the server's finish would use server's
sending state sequence's first value. But let's not digress; the format
stands.

After the handshake:
**Data Message Format (general for any subsequent round):**
struct {
uint32 round_number;
opaque u_value[modulus_length];
// a_value might be omitted if receiver can compute it:



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// but including for completeness
opaque a_value[modulus_length];
uint16 ciphertext_length;
opaque ciphertext[ciphertext_length];
} ASLRecord;
Here, `ciphertext_length` might equal `modulus_length` if a single
field, or could be a multiple if the plaintext was chunked into multiple
modulus-size blocks (like if plaintext is very large, one might do
multiple phi-ns-x transforms sequentially or in parallelthough more
likely you'd just send multiple records).

If we treat each message to be max one modulus worth of data,
ciphertext_length can be fixed to modulus_length except possibly the
last block of a file transfer (but since phi-ns-x isn't a streaming
cipher, better to break each application message separately).

The `round_number` is 32-bit allowing for a large number of messages. In
some contexts (like blockchain), the round might correspond to a block
or tx number.

All multi-byte fields are in network byte order (big-endian) by
convention.

For encoding integers (like p, q, M, u, a):
- They should be in big-endian byte arrays of a fixed length (preferably
the length of M for consistency, padded with leading zeros if needed).
- The prime p or q in server_public_key can be shorter than M (e.g., if
p is 160 bits and M is 256 bits). In that case it can be left-padded
with zeros to some fixed length or encoded with a length prefix. Above
we allow it to be a variable length field with a prefix for length
(`server_public_key<1..+>` means a variable-length vector as per TLS
syntax).

Notably, these structures closely mirror TLS handshake messages albeit
simpler since there's no certificate chain, no separate signature (the
encrypted secret and q^2_masked serve that purpose).

The protocol's normative requirements:
- The client and server MUST verify each other's values as described
(e.g., q^2_masked and handshake decrypt).
- If any verification fails or an unexpected message is received, the
implementation MUST abort the connection.
- All numeric fields are unsigned and non-negative.
- The primes p and q in the key must satisfy all conditions from section
3 (e.g., T composite, etc.) otherwise the protocol's security
assumptions break.

Finally, it's possible to compress some messages in an optimized



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implementation:
- For example, if the cipher suite is fixed and known, ClientHello could
be omitted and the server could start first (like in some PSK modes).
But the full handshake as described is the default.

With the message formats defined, an implementation can
marshal/unmarshal these structures and perform the cryptographic
operations accordingly. This textual specification should be sufficient
to guide interoperable implementations.

5. Example Use Cases

In this section, we discuss how ASL can be applied in various scenarios,
highlighting the benefits and any special considerations for each.

5.1. HTTPS Without TLS (Web Security)

ASL can be used as a drop-in replacement for TLS to secure HTTP
connections (hence "HTTPS without TLS"). In this use case, a web server
would have an ASL server key (Phi-Ns based) instead of an X.509
certificate with an RSA/ECDSA key. The browser (client) would either
have the server's public key (p or q and q^2_masked) pre-loaded (through
some out-of-band means like DNS records, .well-known endpoints, or built
into the application) or would use a trust-on-first-use approach
(similar to how SSH handles host keys).

**Establishment**: When the browser connects to the server, it performs
the ASL handshake:
- The server proves its identity by its ability to decrypt the
ClientEncryptedSecret and by having a q^2_masked matching a published
value. No certificate chain is needed; the trust is directly in the
server's public key. This simplifies the handshake (fewer round trips
than a full TLS handshake with certificate verification) and avoids
reliance on certificate authorities.
- The handshake, as designed, can be as low as 2 round trips
(ClientHello -> ServerHello -> ClientSecret -> ServerFinish), similar to
TLS 1.2 with no client auth. With some optimization (or early data
concept), it could even be reduced, but let's assume 2 RT for fairness
with TLS 1.3 which uses 1 RT but needs certificate in that RT.

**During Data Transfer**: The HTTP requests and responses are sent
encrypted as ASL records:
- The client would send a GET request as plaintext into an ASLRecord
payload. The server decrypts it, processes it, and sends back the
response in one or more ASLRecords.
- The stateful nature of ASL means each request/response increment the
state. In HTTP (which is mostly request-response sequential on a given
connection, at least in HTTP/1.1 or HTTP/2 streams), this aligns well.



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Each HTTP message corresponds to one or more ASL rounds. If pipelining
or multiplexing (like HTTP/2), separate logical streams could either use
separate ASL sequences or be serialized within one sequence - a
multiplexer would handle ordering. Simpler approach: one ASL session per
HTTP/2 stream to avoid head-of-line blocking; but that might be
overkill. It's easier to treat the whole connection as one sequence and
trust ordering at transport.

**Benefits**:
- No reliance on specific cryptosystems that might be broken by quantum
computers (RSA/ECDH). ASL is built on a new assumption believed to be
quantum-resistant.
- Compact public keys: A Phi-Ns key with, say, 192-bit p and q plus
structure can be far smaller than an RSA 2048-bit modulus. It could be
comparable in size to ECC keys, but with PQ security.
- The handshake does not transmit certificates, meaning less data
transferred (helpful for IoT or constrained networks), and avoids the
need for verifying certificate chains (saving computational cost on
clients).
- Privacy: Because no certificates are sent, there's less risk of an
eavesdropper collecting certificate info to track servers. The server's
public key might be well-known via other means, but at least it's not
transmitting a whole identity chain every time.
- Implicit authentication: If the public key is obtained through DNSSEC
or some secure channel, the user agent can authenticate the site
similarly to how it would with a certificate.

**Considerations**:
- Bootstrapping trust: Without the CA system, how does the client know
the server's key is legitimate? This could be solved by a one-time CA-
like attestation (maybe the domain publishes its ASL key in a
certificate-like format signed by a CA, or uses DANE via DNSSEC). This
is a deployment rather than protocol issue.
- Performance: The encryption and decryption operations (involving
modular arithmetic and possibly searching through quadratic solutions)
must be efficient. For web traffic, throughput is important. Profile M
or L from Phi-Ns is appropriate (128-256 bit primes, recursion if
needed). Those should be efficient on modern CPUs. Preliminary estimates
show Phi-Ns operations can be done in microseconds, comparable to RSA or
DH in many cases, but this should be validated.
- Session resumption: As with TLS, one might want to resume a session
without full handshake (to avoid repeating heavy ops). ASL could
implement resumption by having the server issue a session ticket
(encrypted blob containing the shared secrets or a reference) that the
client can present next time. Since ASL is symmetric once established,
resumption is straightforward: it's essentially storing step_u and
current state etc. Implementation of session resumption is an added
layer; the base protocol doesnt preclude it.



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Overall, ASL could provide a secure channel for HTTPS traffic with
potentially simpler mechanics and post-quantum security built-in. The
normative behavior in this use case is exactly as described by the
protocol; the main difference is in how keys are distributed and
managed.

5.2. Blockchain Transactions

In some blockchain systems (particularly permissioned or consortium
chains), there is a need to encrypt transaction payloads on-chain such
that only authorized participants can read them, while others can still
validate some aspects (like knowing it was a valid encryption under a
known key and state).

ASL can be adapted to this scenario, as was partially demonstrated in
the "Rapport" simulation:
- The blockchain network can have a static public key associated with
the chain or with each participant. For example, each participant (node)
might have an ASL key pair. If a transaction is intended for a specific
group, those members share an ASL channel (with shared secret).
- The "Genesis block" or chain configuration would publish relevant
parameters: M, g, and perhaps a common key if one key pair is used
system-wide for encryption (or each participant's public keys if
multiple).
- In the simulation, they had: `Genesis publishes M, g, p_public,
q_anchor, q2_masked + star projections`. This is essentially the network
announcing, "here's the public key (p_public) and associated data you'll
need to encrypt transactions to us; trust is established by q2_masked
etc."
- Each transaction included `(u_k, a_k, C_final)`:
  * `u_k, a_k` correspond to the public state of that round of
encryption.
  * `C_final` is the ciphertext (potentially using an Eq6-like
projection as shown: `C_final = (X + q)^2 - T * a_k^2` in their logs).
- Miners/validators who are not privy to the secret cannot decrypt the
payload, but they can verify that the transaction is well-formed in
terms of state linkage:
  * They can check that the `u_k, a_k` values are correctly derived (a_k
= (u_k + inv(u_k))/2 mod M, which anyone can do).
  * They can check that `u_k, a_k` follow the previous state's
progression if the chain enforces an order (the logs show each tx had a
round number and they likely ensured u_{k} was derived from u_{k-1} by
the shared step).
  * They can store the new state (so that the next transaction from that
channel can be verified against it).
  * If `C_final` uses a construction like Eq6, perhaps the validators
can verify that for the given `C_final` and known public values
(q^2_masked might help here), there's a consistency: actually, without



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knowing T, they can't directly verify (X + q)^2 piece. But if they trust
that any invalid ciphertext would not be producible without secrets,
they might just assume a valid-looking state means it's fine.
Alternatively, they might require the sender to later reveal plaintext
and then verify by re-encrypting (commit-reveal scheme, which was
mentioned in the notes).
- The participants who hold the secret (the ones authorized to decrypt)
will retrieve the transaction, use their private key to decrypt and read
the message. They can later reveal the plaintext if needed for dispute
resolution, and anyone can recompute `C_final` to confirm it matches (as
shown by the log lines: they did a REVEAL verify by recomputing C_final
from plaintext and it matched).

ASL fits this because:
- It provides a way to have evolving state (so you don't reuse the same
key for every transaction, which is good for security and
unlinkability).
- Only those with secret (p, q, etc.) can create valid ciphertexts,
which means unauthorized parties cannot forge a transaction that would
be accepted into the chain (unless they somehow guess the correct state
and produce a valid encryption by chance, which is astronomically
unlikely). This gives implicit access control at the encryption level.
- If a participant leaves or is to be excluded, the key can be rotated
(generate new p, q or just re-randomize abR to effectively create a new
key that yields a new q2_masked, etc.). Future transactions would then
use the new key, and old ones remain encrypted under the old key which
the leaving participant may still have, but new data they can't access.
This is more of an operational procedure.

For a public blockchain scenario where anyone can be a reader, ASL isn't
needed because data is public. But for consortium where privacy is
needed, ASL could run on top of the blockchain as an application layer
or integrated with consensus so that only intended parties see content.

One must also consider performance: blockchain blocks may contain many
transactions, so encryption and decryption should be efficient. ASLs
small key sizes and avoidance of heavy operations per message (just
modular multiplications and one inversion per layer) are beneficial.
Also, the on-chain footprint of each encrypted tx is just a few numbers
(u, a, C) which are all reasonably small (the size of M, e.g., 32 bytes
each for 256-bit security). That's much smaller than storing, say, an
RSA encrypted blob of symmetric key plus the cipher.

The demonstration logs correspond to an ASL channel where multiple
parties (Alice, Bob, Carol) were sharing state (they all had the keys to
update u and could each encrypt messages). That shows ASL supports
multi-sender scenarios if they coordinate state (like in a group, you
need to ensure only one uses a given round or they need to coordinate



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incrementing state). Perhaps in that demo they all share step and doors,
and they just pick next round sequentially when they send a message.

In implementation, to avoid collisions, a group might require a leader
to assign round numbers or have each include the round in transaction
(which they did) and miners ensure no duplicates and order by that.

**Benefits**:
- **Confidentiality**: Sensitive transaction data (like financial
details, private messages, etc. in a block) remain confidential to
authorized parties. The chain still provides integrity and order.
- **Authenticity**: Only someone with the groups shared secret can
produce a valid ciphertext for the next state, preventing outsiders from
injecting transactions that would later decrypt to something. It acts
like a group signature in some way (only group members can "sign" valid
ciphertexts).
- **Post-quantum group security**: Many current blockchain privacy
techniques use ECC (for addresses or for encryption in certain layer-2
solutions). ASL offers a different hardness assumption which might be
safer if ECC is compromised by quantum computers.

**Considerations**:
- Key distribution: Initially, how do group members get the secret key?
That's outside ASL's scope (could be done in person, or using another
secure channel).
- State synchronization: In a distributed setting, ensuring everyone has
the same current state is important. The blockchain itself helps because
the latest block contains the last transaction and thus the last state
(so even if a node was offline, by reading the chain they can catch up
on u_k).
- If the chain allows parallel channels (like multiple disjoint sets of
participants each with their own state), each such channel would have
its own ASL context identified by something like a channel ID, and state
tracked per channel. This is achievable as well.

5.3. Implicit Client Identity Authentication

Traditional protocols often authenticate a client by an explicit
credential (username/password, token, or client certificate). With ASL,
especially in a closed ecosystem (corporate network, IoT deployment), a
client's identity can be tied to possession of a secret key, and the
handshake itself reveals which key (and thus which identity) is being
used, without needing a separate authentication step.

For example, consider an enterprise where each device or user is issued
a unique ASL key pair (or at least a unique secret derived from a master
key). The server (or gateway) maintains a list of authorized public keys
(or their commitments). When a client initiates ASL:



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- The client might include an identifier of which key it's using (or
this could be implicit if the server can try each known public key to
see which matches the q^2_masked provided, but it's easier if client
sends an ID).
- The server receives the handshake with, say, clients p_public and
q^2_masked. If it finds a match in its database, it knows which client
this is. It then proceeds to complete handshake and now all messages are
encrypted.
- If an attacker attempts to impersonate a client, they won't have the
client's secret key, and thus cannot complete the handshake properly
(they'd fail to decrypt or fail to prove knowledge of q by producing
correct q^2_masked).
- Thus, possession of the key is the identity proof. No password or
token needed at session time.

This "implicit identity" is analogous to how SSH key-based auth works:
if you can produce the correct signature with your private key, you are
that user. Here, if you can produce a valid handshake with the key, you
are the authorized device.

Another angle: Suppose the servers ASL key is public for all, but the
shared secret derived in handshake is based on something the client
knows. Actually, that is more like password-authenticated key exchange,
which ASL by itself isn't doing. However, we could incorporate a client
secret in handshake by having the clients encrypted message depend on a
password somehow (e.g., plaintext includes a password that the server
knows and can verify after decrypting K_s). But that is a custom
protocol on top of ASL handshake, not inherent.

Better to just give each client a key:
- It could be a full Phi-Ns key pair or perhaps just a symmetric key
pre-shared to derive something for handshake. For instance, the client
and server could share a long-term secret and use that to generate the
handshake's K_s (like a PSK mode). The client could encrypt something
that includes evidence of the PSK. However, describing that fully is
beyond scope; suffice to say ASL could be extended for PSK (pre-shared
key) mode by mixing a known secret into K_s derivation or verifying a
tag.

In pure public-key mode:
- The server might not want to store all clients' private keys
obviously, it just stores their public half and maybe q^2_masked to
verify integrity.
- Each client effectively has a "certificate" by virtue of their public
key being recognized. The difference is they don't need a separate
signing step; the act of handshake doubles as the signing because only
the real client could produce that encrypted pre-master secret that
results in a meaningful shared key.



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**Use cases**:
- **IoT**: Each device has an ASL key burned in. The central server uses
ASL to communicate, and it immediately knows which device it is talking
to by the key. If a device is compromised, the admin can remove its
public key from allowed list (revocation).
- **VPN**: Users could authenticate to a VPN server by ASL key (instead
of username/password). The connection establishment (like IKE or TLS in
OpenVPN) could be replaced by ASL handshake using the user's key. If
successful, the server associates that session with the user identity of
that key. This is certificate-like but without X.509 overhead.
- **Client certificates alternative**: In web, mutual TLS is rarely used
because it's cumbersome. With ASL, if a client has a key, they can prove
identity seamlessly as part of the handshake without complicated
certificate UI. Of course, the client still needs to have the key pair
available (maybe in OS keystore).

**Security**:
- The strength is that the client key is as hard to forge as breaking
Phi-Ns. There's no online dictionary attack vector since there's no
password, it's public-key crypto.
- The server must protect its list of allowed keys and any mapping to
user accounts.

**Privacy**:
- If clients present a persistent public key, that can be used to track
them across sessions. In some contexts that's fine (enterprise), in
others maybe not (web). There are ways to mitigate tracking by rotating
keys or having the server assign ephemeral identities. But that's deeper
into privacy design.
- The handshake itself reveals which key (hence which client) is
connecting (because p or q^2_masked is essentially an identifier). If
anonymity is needed (like in some protocols), one might not want that.
However, this is similar to client certificates anyway.

5.4. Secure Messaging

Consider an end-to-end encrypted messaging application (like a chat
app). Typically, such apps use protocols like Signal, which involve
Diffie-Hellman key exchanges, X3DH, double ratchets, etc., to ensure
forward secrecy and authentication. ASL could serve as an alternative
core for some of those steps, or even the entire session.

Scenario:
- Alice wants to send an encrypted message to Bob. She knows Bob's ASL
public key (obtained from Bob's profile or directory).
- They could perform an ASL handshake to establish a shared secret and
then send messages. But if they are not online at the same time for a
handshake, you could do an offline setup:



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  * Possibly one party can precompute a handshake message and the other
completes it when they come online. This is tricky because ASL as
described is interactive. But maybe Alice can send "ClientHello +
ClientEncryptedSecret" to Bob (through the server) while Bob is offline.
Bob when coming online can process it and send "ServerHello +
ServerFinish"? The ordering would differ because normally server goes
second. Alternatively, they'd just wait until both are online.
- More straightforward, if asynchronous: treat each message as its own
ASL "handshake+data" maybe, or use pre-shared static keys with ratchets
on top.
- However, if interactivity is allowed (both online concurrently, like
voice call or live chat session), they can do a full ASL handshake
directly peer-to-peer and then chat with messages as ASL records.

In such a chat:
- Both peers likely have devices with their key pairs (like each user
has a long-term key).
- Mutual authentication can occur by each verifying the other's
q^2_masked through a fingerprint exchange (like scanning a QR code of
the key fingerprint out-of-band, similar to how Signal verifies
identities).
- Once they have a session, each message is encrypted with evolving
state. This inherently provides **forward secrecy** across messages: an
attacker who compromises a key after some messages were sent cannot
decrypt the earlier messages if those messages involved state that isn't
reproducible from just the static key. Actually, careful: if an attacker
gets p and abR, can they decrypt old messages? They still need the per-
session shared secret K_s or at least the states. If those states were
derived from K_s which was ephemeral and not stored, then yes, forward
secrecy holds. So to ensure FS, one or both must contribute an ephemeral
in handshake (like ephemeral Diffie-Hellman in TLS). ASL handshake as
given uses the static server key and an ephemeral client secret K_s. But
the static server key by itself means if server key is compromised, the
handshake could be decrypted because the attacker can simulate it?
Actually, they'd need the transcript. If someone logs all handshake
messages and later the server key gets stolen, could they derive the
session keys? Possibly yes, because with p, q they could decrypt
ClientEncryptedSecret (X_secret). So without an ephemeral server key, a
long-term key compromise does compromise past sessions. So strictly, the
current ASL handshake is not forward-secret if the server's (or main key
holder's) static key is compromised. Its like RSA TLS: break RSA key,
you can decrypt old pre-master secrets.
- So for messaging, they likely want forward secrecy. That can be
achieved by having both parties use ephemeral keys in an ASL handshake
(like I mentioned, each could generate an ephemeral phi-ns pair,
exchange them in the handshake and derive secrets combined). Or simpler,
just do a fresh handshake with ephemeral keys for each message or each
session, but that might be too slow if frequent.



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- Alternatively, ASL could be integrated into a double ratchet scheme:
for each message, the state update acts like a ratchet step since it
changes the keys. Provided that the state updates are not reversible
without the prior key, then even if one message key is compromised,
previous keys are safe. But if the static key is compromised, all might
be at risk.
- Possibly combine static identity keys (for authentication) and
ephemeral session keys (for encryption). This is akin to how
Axolotl/Signal does it: you have identity keys and ephemeral one-time
keys. One could design an "X3ASL" handshake: exchange static keys
signed, plus ephemeral keys in an ASL style.

The details can get complex. The main point: ASL provides a structure
that can do the heavy lifting of encryption with evolving keys, but one
might still need to incorporate ephemeral keys for full perfect forward
secrecy if desired.

If forward secrecy is less a concern (say within a short live chat
session), and key compromise is unlikely in real-time, ASL as is may be
fine.

**Benefits** for messaging:
- Small messages overhead: each message only needs to carry a small
state and ciphertext. Compare to, e.g., sending a whole new RSA-
encrypted symmetric key per message (which some naive systems do), ASL
is more efficient.
- No need for additional MAC each message if relying on state.
- The same library that does handshake sets up continuous encryption,
simpler architecture.

**Group chats**:
- If multiple parties in a group chat, they can share one ASL channel
(like the blockchain example, but now interactive).
- Alternatively, pairwise channels or using group keys (beyond our
scope).
- ASL's design with hidden factors could be exploited in creative ways
for group: e.g., each group member could hold a piece of R such that
only if they cooperate can they decrypt? That's speculative and not
standard ASL usage.

5.5. Encrypted Tunneling (VPN)

A VPN or any encrypted tunnel (like an IPsec replacement, or a simple
point-to-point tunnel between networks) can use ASL to establish a
secure link.

Typically, IPsec has IKE for key exchange then uses symmetric ciphers
for bulk data. TLS-based VPNs (like OpenVPN) do similar with TLS



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handshake then symmetric cipher (AES) for data.

With ASL, we can unify that: the handshake and the data encryption use
the same primitives and keys:
- The gateway and client perform an ASL handshake to authenticate (the
gateway proves identity via q^2_masked, client perhaps via PSK or
certificate equivalent).
- They derive a shared secret and trajectory parameters.
- Then they start forwarding IP packets or frames through the ASL
channel. Each packet is encapsulated in an ASLRecord.
- Because of stateful encryption, this looks similar to a stream cipher
in usage: packets go in order. If out-of-order delivery is needed (like
multiple parallel packet flows), one might consider independent channels
for each or accept the sequential requirement (which might add latency
if one packet is lost, subsequent can't be decrypted correctly until
maybe that state is recovered or skipped).
- Perhaps treat each packet independently by resetting state if you
don't want head-of-line issues. But that loses the benefit of evolving
state (though you could still do a quick re-handshake or treat each as
separate message with small state changes).

For a reliable transport or for link encryption where ordering is
ensured, ASL is fine.

**Post-quantum benefit**:
- Many current VPNs would need to switch to PQ algorithms; ASL is a
candidate with unique properties.
- As a pure arithmetic algorithm, it might be easier to implement in
embedded routers without special hardware (and not be subject to certain
side-channel issues if careful, though side-channel protection would be
needed as always).
- The overhead per packet is just a few bytes plus the computational
overhead which should be manageable (maybe on the order of a few hundred
multiplications per packet, which on hardware can be optimized).

**Integration**:
- ASL could be a new protocol at either layer 4 (like a custom reliable
protocol carrying encrypted data) or layer 2/3 (like how IPsec ESP
works).
- One could even embed ASL within existing protocols by using reserved
bits (for example, encode u and a in an IP option header - not realistic
publicly, but in a closed system who knows).

**Security**:
- MITM is prevented because the handshake authenticates the server (and
possibly client). Attackers on the path cannot decrypt or inject because
they can't predict the state or produce valid ciphertext.
- Even if one packet is captured, without the key they can't glean



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anything. And if one tries to replay a packet, the state will have moved
on so the receiver would find the round number out of date and drop it,
thus providing replay protection.

**Comparison**:
- Compared to IPsec's use of AES-GCM (which provides confidentiality and
integrity with an IV and counter per packet), ASL provides a somewhat
analogous function: the evolving state is like a counter/IV, but with
secret evolution making it unpredictable to outsiders (whereas in AES-
GCM the counter is usually not secret but just a nonce).
- ASL inherently authenticates the data because forging requires the
key. AES-GCM similarly prevents forgery due to the tag. ASL's approach
might lack an explicit tag, but the non-linearity could serve similarly.
Still, one might consider adding an explicit lightweight tag if
paranoid.
- Throughput: AES is faster with hardware support. ASL is software-
based, but if its operations can be vectorized or offloaded (like
modular multiplication can be done with big integer libraries, maybe not
as fast as AES-NI instructions though). For moderate speeds (e.g., <1
Gbps) ASL should be fine in software on modern CPUs. At very high
speeds, one might explore hardware acceleration (maybe FPGAs could
accelerate mod M operations, or a custom ASL co-processor).
- The benefit is not needing separate systems for key exchange and
encryption - it's unified.

6.
4.x Encryptor Operation and Knowledge Model

This section specifies precisely the role, capabilities, and limitations
of the encrypting party in the Authenticated Secured Layer protocol
using
Phi-NS.

The encryptor does NOT possess the private root p, the value q, nor the
quadratic gap T = q^2  p^2. The encryptor is not required to know any
secret factorization, atom ordering, or private trajectory parameters.

The encryptor operates exclusively on public, evolving state values.

For each encryption round k, the encryptor is provided with a public
state value u_k  Z_M*.

From this public state, the phase value a_k is deterministically derived
as:

    a_k = (u_k + u_k^{-1})  inv2 mod M

where inv2 = 2^{-1} mod M.



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The encryptor then applies the sequential Phi-NS transformation to the
message.

Let the plaintext message be represented as an integer m  Z_M.

The encryption recurrence is defined as follows:

    p_0 = m

    For i = 1 ... L:

        y_i = (p_{i1} + a_k) mod M

        p_i =
            ((y_i^2  y_i + p_{i1}  a_k)  y_i^{-1}) mod M

where y_i^{-1} denotes the modular inverse of y_i in Z_M.

The ciphertext output of the encryption process is defined as:

    C = p_L

Optionally, the ciphertext MAY be further bound to the quadratic
structure by applying a quadratic projection, as defined elsewhere in
this document.

At no point does the encryptor require knowledge of p, q, T, or any
private atom or star structure. The encryption process is entirely
public and deterministic with respect to the public state u_k.

The security of the scheme relies on the infeasibility of reconstructing
the inverse trajectory without possession of the private parameters and
the correct internal ordering of transformations.


Security Considerations

This section analyzes various security aspects of ASL, reiterating some
points from earlier sections in a concise manner, and adding any
additional considerations. ASL is designed to provide confidentiality,
integrity, and authenticity in the presence of active adversaries, and
to resist both classical and quantum cryptanalytic attacks. However, as
with any cryptographic protocol, correct implementation and parameter
choices are critical.

6.1. Man-in-the-Middle (MITM)

ASL explicitly addresses MITM attacks through its handshake design. An



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MITM is an attacker who intercepts communications and attempts to
impersonate each party to the other. In the ASL handshake:
- The server authentication is based on possession of the secret prime
(or its factors). The client verifies the server by checking the
`q^2_masked` value against a trusted reference and by ensuring the
server can decrypt the client's secret message. An attacker in the
middle who does not have the server's private key cannot satisfy both
conditions. They might forward messages to the real server (attempting a
relay attack), but the handshake verification (e.g., the final encrypted
verify message with state binding) will fail unless the MITM lets the
actual server complete the handshake, in which case the MITM isn't
actually decrypting anything (just tunneling).
- If the MITM tries to use their own key pair to talk to the client, the
`q^2_masked` check will fail (client won't recognize this rogue server's
identity as legitimate). If they try to use their key with the real
server, the server will similarly not recognize the client if client
authentication is expected, or the server will complete handshake but
then the MITM can't use the session because the client side didn't
complete it with them properly.
- Thus, as long as the client has a correct trusted copy of the server's
public-key identifier (and the server is not compromised), a MITM cannot
pretend to be the server. This is equivalent to saying ASL, like TLS,
assumes an authenticated server public key distribution. If that initial
distribution is via a secure channel (like a known fingerprint, or DANE,
or a CA-signed object), MITM is thwarted. If it's not (like no
verification, just take whatever comes first), then a MITM during first
contact could trick the client into trusting a fake key (TOFU scenario).
That is not a protocol flaw but a bootstrapping issue.
- For client authentication (if used), a MITM similarly cannot
masquerade as the legitimate client to the server because they cannot
produce a valid handshake under the client's public key if they don't
have the private components. The server will either not accept the
unknown key or, if the MITM tries a known client's ID but doesn't have
key, it will fail at handshake verification.

ASL also binds handshake and record layer tightly via the evolving
state. After handshake, a MITM who didn't fully participate (and succeed
in impersonation, which we argue they can't if checks are done) cannot
inject data later:
- Suppose a MITM recorded some ciphertext and tries to replay or inject
it. If they don't know the current correct state or key, the receiver
will find the `u_k` is not the expected next state (replay detection) or
will fail to decrypt properly. Because the state transition is one-way
and secret-driven, the MITM cannot guess a future valid state to inject
fake ciphertext.
- Meet-in-the-middle attacks in key search (where attacker tries to
split a key into two parts and compromise one part at a time) do not
apply to ASL's core because the key isn't used in a way that halves the



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problem. The entire structure has to be guessed as a whole (see section
3.4 on no partial disclosure).
- A specific MITM concern is protocol downgrade: if there were multiple
cipher suites or versions, an attacker could try to force the use of a
weaker one by dropping messages. ASL's ClientHello/ServerHello version
negotiation should be done with care (like TLS, the client should abort
if server picks something not actually offered). We assume any alternate
modes (like turning off injection or using smaller keys) are either not
present or if present, chosen in a way that a MITM cannot cause
insecurity (by always using secure defaults or requiring client
consent).

In summary, if ASL is executed correctly with authenticated keys, MITM
is effectively mitigated. Implementers MUST ensure to perform the
verification of `q^2_masked` and handshake results. Skipping those
checks or ignoring authentication (for convenience or performance) would
open a door to MITM, just as not checking a TLS certificate would.

6.2. Collision Resistance

"Collisions" in this context can mean a few things:
- Two different keys yielding the same public parameters (commit
collisions).
- Two different plaintexts yielding the same ciphertext under different
keys or states (encryption collisions).
- Collisions in the sense of cryptographic hash (if we used any).

**Key/Public Parameter Collisions**:
- Could there be two distinct secret tuples (p, a, b, R) and (p', a',
b', R') that produce the same public outputs (either same q or same
commit)?
- Having the same q would mean either p = p' or a scenario like p and p'
both associated with the same q (which from q^2 - p^2 = q^2 - p'^2
implies p = p', so either p = p' or p = -p' mod something but since
these are integers, basically p equal or just sign difference irrelevant
here because primes are positive).
- So no two different p values should yield the same q.
- Could two different abR structures yield the same serialized commit?
The design of serialization and commitment (assuming an ideal hash or a
well-chosen masked scheme) should make collisions computationally
infeasible. We rely on that as a design assumption: e.g., if q^2_masked
is chosen as say q^2 mod N for some number N, an attacker might try to
find a different q' that yields same q^2 mod N, but that would require
solving q^2  q'^2 (mod N), which often has the trivial solutions q'  q
(mod N). If N is large (like a 512-bit number) and not specially
structured, the chance of a random different q colliding is negligible.
If an attacker found a colliding key that produces the same masked
commit as a target, they could impersonate that target's identity. To



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guard against that, one should choose the commit function to be
collision-resistant (like a cryptographic hash or sufficiently large
modulus to make solving that congruence as hard as guessing q).
- We consider that aspect in design: it's RECOMMENDED to use a proven
collision-resistant method for the commitment. Using a hash like SHA-256
of the serialized abR would be safest (practically no collisions). Our
avoidance of SHA in core was philosophical, but one might reintroduce a
hash just for identity safety. If not, the masked approach used must be
analyzed for collision risk.
- In practice, the chance of two legitimate users picking keys that
collide in commit is astronomically low if done right. The bigger risk
is an attacker deliberately computing a colliding key pair if any
weakness is found in commit. That's a niche attack; presumably as hard
as breaking the scheme itself if commit is well-designed.

**Ciphertext Collisions**:
- Could two different plaintexts produce the same ciphertext under the
same key and state? Ideally not, as that would violate invertibility for
the legitimate decryptor (they wouldn't know which plaintext was
intended). Given phi-ns-x is a permutation on Z_M for fixed a_k (except
the small chance of hitting the y_i = 0 case), it should be one-to-one
(maybe not strictly a permutation on the entire space if y=0 is
excluded, but on the domain of allowed inputs it should be injective).
- The transform looks invertible if you have the secret; it's likely
that for each possible X there is exactly one plaintext that would
produce it (or at most a small constant number given the quadratic
nature, but the decryptor can figure out the correct one via key).
- So collisions in this sense are not expected; each encryption mapping
is unique per message. There's no known smaller subspace that two
different messages would converge to the same X because that would imply
solving p_L = constant for two different p_0 given the parameters  a
difficult equation with no reason to have two solutions except trivial
symmetric cases, which the key structure would break.
- If optional masking like Eq6 is used, that formula is still injective
given secret info; but an attacker might see multiple ciphertexts that
numerically collide (e.g., if a certain pattern made two different X
yield same C, they'd be indistinguishable to attacker, but decryptor
would not confuse them because they'd decrypt differently). However, one
must ensure that doesn't lead to a chosen-ciphertext vulnerability
(shouldn't, because attacker can't control X without key anyway).
- Also, sequence numbers (round) effectively make even identical
plaintext encrypted at different times yield different ciphertext
because a_k (hence the transform) changes each round. This provides
semantic security (no pattern leak if same message repeated, unlike a
deterministic cipher).

**Hash Collisions**:
- If any hashes (like for commit or handshake verify) are used,



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collision resistance of those is required to avoid impersonation or
tampering. Standard hashes like SHA-256 are fine. If an attacker found a
hash collision in, say, the commit function, they could potentially
produce a different abR that yields same commit, undermining key
authenticity. This is theoretical (depending on commit scheme).
- Or if using hash for handshake transcript, collisions are not a big
worry unless an attacker could craft two different transcripts that
produce same hash and trick one party, but that seems remote and anyway
if handshake was manipulated that far, likely other parts would break.

Overall, collisions are not seen as a practical threat in ASL when
parameters and algorithms are chosen appropriately. Nevertheless:
- **MUST**: The commit function for key identity must be collision-
resistant (to at least 128-bit security). Using a secure hash function
is recommended unless a purely arithmetic method with equivalent
hardness is in place.
- **MUST**: The encryption function should never be intentionally used
in a way that non-injective behavior could occur. For instance, avoid
using an M that is too small such that the plaintext space is much
larger than M (which could cause wrap-around collisions).
- **MUST**: Each encryption round use a unique state to avoid trivial
collisions (like encrypting two messages with same state would produce
same ciphertext if plaintexts equal; but state reuse must be prevented
like IV reuse in any cipher).

6.3. Secret Leakage and Oracle-Free Design

One of the core security principles of Phi-Ns and ASL is the absence of
any "oracle" that leaks information about the secret key or intermediate
states:
- During handshake, the server does not disclose anything that
correlates directly with its secret. The `q^2_masked` value is a one-way
function of q. The encryption/decryption process either succeeds or
fails but does not reveal partial info: if an attacker guesses a wrong p
and tries to decrypt something, they either get a wrong result or no
result, with no indication of how far off they were.
- The protocol does not use branch conditions or different responses
that depend on secret values (at the spec level). Implementations should
ensure to follow this: e.g., a naive implementation might accidentally
take a different amount of time to decrypt based on how factors align,
or return distinct error messages for different failure reasons. That
could leak info via side-channels or behavior. An implementation
**SHOULD** use constant-time arithmetic where feasible and treat all
failures uniformly (just "handshake failed" without saying why).
- In the data phase, an attacker cannot query the decryptor with
arbitrary ciphertexts to get hints. If they send a malformed ciphertext,
the receiver will drop it or potentially respond with a generic error or
de-sync. There is no decrypt-oracle giving success/failure on a per-



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packet basis in a useful way; at best, an attacker could flood with
random ciphertexts and see if the receiver at some point processes one
(meaning it was valid). But the odds of guessing a valid ciphertext and
state is essentially zero. The receiver might also have anti-replay and
anti-desynchronization logic (like not resetting state on a bad message,
etc.) to avoid giving even timing clues. In short, forging a valid
ciphertext is as hard as breaking the crypto, so no oracle helps.
- The public values u_k and a_k leak no secret info by design. `a_k` is
related to `u_k` by a simple formula with no secret. `u_k` evolves by
multiplication with g^step_u. Now, one might worry: if an attacker sees
two successive u values, can they solve for step_u? This is essentially
solving g^step_u = u_{k+1} * inv(u_k) mod M. That's a discrete log
problem. If M is large enough (e.g., 2048-bit prime), discrete log is
infeasible classically and also by known sub-exponential algorithms. But
note: discrete log is not NP-hard in general, it's sub-exponential (like
128-bit security might require 3072-bit modulus). But also note: a
quantum attacker could solve discrete log mod M in polynomial time
(Shor's algorithm) if they had a quantum computer.
- Does that mean if someone had a quantum computer, could they find
step_u and then predict future states? Yes, if M was a standard prime
field, Shor would find step_u from a single state pair. That would let
them predict all future u/a. But would that let them decrypt content?
Not directly, because they still don't have p or the door values. They
could mimic the state sequence and attempt to encrypt their own messages
or maybe mount some known plaintext attack if they guess content.
However, just knowing state might help them attempt to brute force p by
having more structured equations. But arguably, even with known u
sequence, the core phi-ns problem remains (they have q possibly if
server is PK-q, they have p if PK-p). Actually, if PK-q mode and
attacker gets step from states, they know how u evolves. They could then
generate u and a for future, but still can't decrypt without p. They
could however attempt to send messages (since encryptor only needs a and
public anchor q). Uh oh: if attacker knows step_u and has q (which is
public in PK-q), they could now encrypt messages that the server would
accept because they can predict the state and they know q. That breaks
authenticity in PK-q mode if step is exposed. This is why leaking step_u
is dangerous. So we must rely on discrete log hardness for adversaries.
- Therefore, to avoid that, one might choose M to be a safe prime and g
a generator, making discrete log as hard as possible for given size. For
PQ safety, though, discrete log is vulnerable to quantum. But the idea
is the attacker still couldn't decrypt past messages, but they could
inject future ones if they break step. This suggests an interesting
point: in a post-quantum scenario, even if phi-ns core is safe from
quantum (no known algorithm to recover p from q easily), the use of a
standard finite field for state might be a weakness because of Shor's
algorithm. If a quantum adversary can derive step_u quickly, they can
impersonate a client or server in sending data (though they'd still need
to forge the proper door values if used).



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- Potential mitigation: one could use a different group where discrete
log is also thought PQ-hard (like a large non-abelian group or some
lattice structure) for the state sequence. That becomes complex.
Alternatively, ensure door values (stars) are used, because even if
attacker predicts a and tries to encrypt, if they don't know b_i
sequence, their ciphertext will be wrong. So the secret injection saves
the day: it means even with knowledge of step and q, attacker can't
create valid ciphertext because they'd miss the b_i.
- So it's important: **If quantum adversary is a concern, use door
injection or another secret element in encryption beyond just a and q**.
This way, predicting the public state isn't enough to forge messages.
- Also, after quantum, discrete log mod a large classical prime might
become easy, so step would be known. However, recovering p from q (the
phi-ns problem) might still be hard (super-exponential). So
confidentiality might hold, but authenticity could suffer without the
extra secret in each encryption. Therefore, for long-term PQ robustness,
it's RECOMMENDED to enable the layer injection feature with secret
values (stars/doors) when using ASL if one expects adversaries with
quantum capabilities that could target the state update.

**Side-channel leaks**:
- While not explicitly in protocol, implementers must watch for timing,
cache, and other side-channels. For example, big integer operations may
vary in time based on operands (leading zeros, etc.), which can
sometimes leak bits of secrets. If a side-channel attacker on the same
machine could observe decryption time variations, they might glean info
about p or R. It's a complex analysis for phi-ns, but general best
practice: use constant-time arithmetic for key ops, avoid secret-
dependent branches or memory access.
- Storing of secrets: The private key (p, a, b, R, and seeds) should be
protected in memory. After use, if a session key K_s is derived, it
should be zeroed out after use or when session ends. Similarly, if
recursion adds additional secrets, treat them carefully.
- Randomness: The handshake relies on random `K_s`. If the clients
random generation were poor, an attacker might guess K_s and break that
session. Also, the servers seed for key generation must be truly random,
else the key structure might have hidden correlations. So a robust
CSPRNG is needed for keys and nonces.
- The protocol has no dependence on any external random oracle at
runtime besides those nonces and ephemeral secrets.

In summary, ASL's design avoids deliberate information leakage.
Implementers must maintain that property by following cryptographic
coding standards. The protocol is safe against chosen-message attacks
since an attacker cannot cause the legitimate parties to output
something that reveals the key (there's no encryption oracle giving out
something like plaintext bits or error messages keyed on secret). The
worst an attacker can do is send junk and get disconnected.



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6.4. Post-Quantum Resilience

The cryptographic strength of ASL in a post-quantum setting is a key
motivation for its design:
- The Phi-Ns problem (recovering p given q and commit) has no known
quantum sub-exponential attack. Grover's algorithm gives at best a sqrt
speedup to brute force, which is negligible if the key space is super-
exponential. The abR structure can be made large enough that even a
quadratic reduction is insufficient.
- To contrast: RSA or ECC fall to Shor's algorithm easily. Lattice-based
schemes have better PQ standing but often larger keys and uncertain new
problems. Phi-Ns offers another independent hardness assumption.
- That said, one must consider *all* parts of the protocol under quantum
threat:
  * As mentioned, the discrete log in Z_M could be a weak link (since
thats traditional math). If a sufficiently large quantum computer
existed and M was standard size (like 2048 bits), an attacker could find
step_u, step_star etc. If they also have q (in PK-q mode), they could
then attempt to impersonate as discussed. If in PK-p mode (q secret),
knowing step doesn't directly give them ability to decrypt or send,
since they still lack q and the injection secrets. So PK-p plus
injection seems robust even against that quantum capability, because
they'd still face the main phi-ns puzzle or have to solve for q.
  * If for some reason we used a known hash function for commit (like
SHA-256), note that Grover could find preimages in 2^128 rather than
2^256, which is still fine. Collision-finding doesn't benefit as much
(quantum collision algorithms get 2^(n/3) complexity for n-bit hash, so
SHA-256 would have 2^85, which is borderline, but then again one could
switch to SHA-384 or SHA-512 in future or rely on the difficulty of
forging a full key that matches a commit).
- So, in a scenario where both sides anticipate quantum eavesdroppers,
some adjustments can strengthen ASL:
  * Use larger M (size such that even quantum discrete log is hard;
unfortunately, quantum DLP is polynomial, so any finite field is
breakable at large scale).
  * Or use double encryption: one could combine phi-ns-x with a one-time
pad or symmetric cipher as a fail-safe. For instance, after handshake,
derive a symmetric key and use a quantum-resistant symmetric cipher
(like AES) to encrypt data, and use phi-ns-x mainly for key exchange and
authentication. This hybridity might be overkill but it's a path.
  * More straightforward: rely on the fact that to break ASL fully, the
attacker needs to break phi-ns (to get the actual content) or at least
the door secrets to forge data. We assume phi-ns stands firm against
quantum, so confidentiality remains. For authenticity, ensure some
secret piece remains unknown even if states are known. That we have done
via door values.

We emphasize that as of writing, quantum computers large enough to



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attack 2048-bit discrete log do not exist, but they might within a
couple decades. Meanwhile, phi-ns claims resistance due to no known
algorithm. It's newer, so it warrants scrutiny from cryptographers to
solidify that claim.

One also should consider **data longevity**: If someone records ASL-
encrypted traffic now and stores it, and in 20 years gets a quantum
computer, could they decrypt it?
- They would have q (if PK-q) or p (if PK-p) and M, etc. They might
break the state progression (trivial with quantum as said) and then
they'd know all a_k, but still they'd face the challenge of deciphering
each message without p or q. If by then algorithms to invert phi-ns
appear (quantum or even classical improvements), then confidentiality
would be lost. So, if extremely long-term secrecy (decades) is needed,
one might prefer to combine ASL with another PQ algorithm (hybrid
encryption) just in case. However, if one believes in the strength of
the combinatorial assumption, storing exabytes of classical cipher data
might remain safe.

**One more angle: lattice or code-based integration**: Could phi-ns keys
be used in a lattice-like structure to update state instead of discrete
exponent? Possibly future research. For now, we keep to classical groups
but note the caveat.

In summary, ASL is designed to be post-quantum secure to the best of
current knowledge, but like all schemes, that depends on unproven
hardness assumptions. We have identified potential quantum weaknesses
(like state DLP) and provided mitigations (like secret injection). It is
the responsibility of implementers targeting PQ security to enable those
features and choose parameters accordingly (e.g., at least Profile L:
256-bit p, q and perhaps deeper recursion or bigger R to frustrate any
unknown quantum strategies).

6.5. Unpredictability of States

A crucial security aspect of ASL is that the public values broadcast
with each message (the state: u_k and a_k) do not allow an adversary to
predict future states or deduce past secrets. This unpredictability
ensures **forward secrecy of state** and also contributes to
authenticity (an attacker can't inject a valid future message because
they can't guess the correct state and ciphertext combo).

- The sequence u_k is essentially a pseudo-random sequence from the
attackers viewpoint. Without knowing step_u, the best the attacker can
do is see it as some random elements in Z_M. Even if they collect many
pairs (u_k, u_{k+1}), it doesn't help them extrapolate unless they solve
the discrete log problem to get step_u. As discussed, classically that's
essentially impossible for large M, and even quantumly it's the one weak



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link but mitigated with injection or an alternate approach.
- Because of this, the **probability of an attacker correctly guessing
the next valid u_{k+1}, a_{k+1} pair is 1/(M * something)** basically
extremely small (like picking a random 2048-bit number correctly).
- Likewise, if an attacker tries to replay an old message, the state
(round number or u value) will likely not match the expected next state,
and the receiver will detect that. So state sequence provides a built-in
anti-replay mechanism. To be safe, implementations should:
  * Track the last seen round or state from each sender.
  * Reject any message that does not have round > last_round (or if
round numbers reset, use another mechanism, perhaps window).
  * If round numbers are not explicit and just inferred by order, then
any out-of-order arrival should be either ignored or cause a reset.
Using explicit numbering is safer to detect replays.

- **Non-deducibility of states** also means even if some state gets
revealed (imagine an insider or side-channel reveals the current u_k),
the attacker still cannot compute backwards to prior states without step
(discrete log again). So backward secrecy is also there: observing a
current state doesn't tell you what previous states were used (although
an attacker might have seen them anyway if they are public, but if they
missed some, they can't derive them from later ones easily).
- If an attacker could somehow find a pattern or relation in the
sequence (like a linear relation), that would break unpredictability.
But ASL's usage of multiplicative group means it's essentially a power
sequence (u_k = u_0 * g^{k*step_u}). This is exponential progression mod
M, known to be a sequence with maximal period if g is generator and
step_u not a multiple of the order. It's as unpredictible as discrete
log assumption. There's no linear relation like LFSRs. So it's solid.
- The inclusion of optional star rotation can make the state not just
one sequence but a more complex one: if we modulate or sometimes replace
u progression with another factor injection, to an external observer the
sequence might look even more random (not just a single exponentiation
pattern). However, from a pure unpredictability perspective, a single
exponent sequence is already cryptographically random if DLP is hard.

- **No partial prediction**: If an attacker sees u_k and u_{k+2} but
misses u_{k+1}, can they compute it? That would require solving
something like u_{k+2}/u_k = g^{2*step_u}, thus g^{step_u} would be a
square root of that ratio, which is another discrete log essentially
(two-step gap doesn't help; it might even give them a slight advantage
like they know g^{2*step} now; in some cases they could attempt to take
square root if group structure allowed multiple solutions, but since mod
prime, taking sqrt is also non-trivial relative to DLP if step is odd
etc).
- It's theoretically possible if step_u was small, an attacker could
brute force it by checking successive powers. Therefore, step_u should
be a large random integer mod (M-1). If one naively set step_u to 1 or a



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small number, an attacker could just try those. The specification says
step_u is a shared secret and not published, but we should also ensure
it's not too small or guessable. Essentially treat step_u as ~|M| bits
of entropy.
- Also if M-1 has small factors and step fell into a subgroup of small
order, u might cycle in a shorter loop. That's an edge case: one should
choose M (like a safe prime where M-1 = 2q, q prime, or at least with a
large prime factor) and choose step not a multiple of any small factors
to ensure full period. Possibly define that step must result in u having
large order (maybe even exactly the largest prime factor of M-1).
- If u cycles quickly, predictability might not be easier (still need
log, but repetition could allow some statistical detection). Anyway, use
large subgroup if not full group.

- During an ongoing session, compromise of secrets provides at most
future state knowledge but cannot retroactively decrypt earlier messages
due to forward secrecy of ephemeral secrets, if ephemeral were used.
Without ephemeral, compromising static key allows decrypting recorded
past sessions.

Another aspect: if an attacker can't predict states, they also can't
perform any kind of chosen-plaintext correlation attack. For example, in
some stream ciphers, if state/IV reuse happens, attackers exploit that.
Here, state reuse should never happen as long as step isn't zero mod
order or such. There is no key reuse per message because it's always
evolving.

Summing up:
- The state sequence must be treated as a critical part of security.
It's public but pseudo-random. Breaking its unpredictability reduces ASL
to a simpler form that might be attackable.
- Implementers must use strong parameters for the state evolution to
maintain unpredictability. They must also enforce the protocol rules (no
state reuse, detect replays).
- If those are in place, an adversary gains nothing useful from
observing states besides synchronization info.

One might wonder if the public nature of a and u could ever leak info
about the secret key (p or q). According to the design, they should not:
p and q are only present in the encryption transform internally on
receiver side. The values u and a are independent of p (they depend on
step and g, which are unrelated to p; step is derived from K_s which
might indirectly be influenced by p only if something like we derive
step from private structure in a fancy implementation, but typically
step is handshake ephemeral or from K_s).
If an implementer foolishly derived step_u directly from p or R (like
making step = some function of secret key), then u sequence could leak
something about the key if someone solved DLP. That is not recommended.



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step should be truly random from handshake secret, not deterministic
from key. Keep the key and state mechanism separate except the key is
needed to decrypt.

This ends the security analysis. Overall, ASL appears to achieve its
goals under standard and new hardness assumptions. Nevertheless, it
should undergo extensive analysis by the community. As any new
cryptosystems, it requires scrutiny, and while we have outlined reasons
for its strength, only time and effort will validate them. Users of ASL
in critical systems should follow developments and be ready to adapt
parameters or schemes if any weaknesses are found.

7. IANA Considerations

This document does not require any actions from IANA. No new protocol
numbers, ports, or registries are defined here.

(If ASL were to be assigned a standard port or content type, that could
be addressed in a future document. As of now, it operates at an
application or session layer with no centralized coordination needed.)

8. References

8.1. Normative References

[RFC2119]  Bradner, S., "Key words for use in RFCs to Indicate
Requirement Levels",
           BCP 14, RFC 2119, March 1997.

[RFC8174]  Leiba, B., "Ambiguity of Uppercase vs Lowercase in RFC 2119
Key Words",
           BCP 14, RFC 8174, May 2017.

[PhiNs]    Nsiangani, P. J., "Phi-Ns Cryptosystem: Structured Quadratic
Gap Hardness for Public-Key Cryptography",
           draft-nsiangani-phi-ns-01 (work in progress), February 2026.

8. References

8.1. Normative References

[RFC2119]  Bradner, S., "Key words for use in RFCs to Indicate
           Requirement Levels", BCP 14, RFC 2119, March 1997.

[RFC8174]  Leiba, B., "Ambiguity of Uppercase vs Lowercase in RFC 2119
           Key Words", BCP 14, RFC 8174, May 2017.





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8.2. Informative References


[NISTPQC]  National Institute of Standards and Technology, "Post-Quantum
           Cryptography Standardization Project", 2016-2025.

[ANSSI]    Agence Nationale de la Securite des Systemes d'Information,
           "RGS v2.0  Rfrentiel Gnral de Scurit", 2020.

[NCSC]     National Cyber Security Centre, "Cryptographic
           Recommendations for TLS, VPNs, and Messaging", 2022.

[Lamport]  Lamport, L., "Constructing Digital Signatures from a One Way
           Function", October 1979.

[SHA256]   Dobbertin, H., Bosselaers, A., and Preneel, B., "The Hash
           Functions MD5, SHA-1 and RIPEMD", 1996.

[PhiNs]    Nsangani, J., "Phi-Ns: Quadratic Structured Decomposition
           for Asymmetric Cryptography", Patent Pending, 2025.



Author's Address

   Parfait Junior Nsiangani (Editor)
   Email: jnsiangani@gmail.com

             Figure 1: Original draft content (imported)

2.  References

Author's Address

   P.J. Nsiangani (editor)
   ACE Working Group















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name: draft-ietf-lamps-est-renewal-info
rev: 00
title: Certificate Renewal Recommendations for Enrollment over Secure Transport
time: 2026-02-12T17:10:42Z
group_uri: /api/v1/group/group/2164/
expires: 2026-08-16T17:00:28Z
states: ["/api/v1/doc/state/1/", "/api/v1/doc/state/150/", "/api/v1/doc/state/38/"]
categories: ["Data formats/interop"]
tags: []
fetched_at: 2026-02-27T22:56:09.014965+00:00
abstract:
This document describes an extension to RFC7030, Enrollment over
   Secure Transport to give an indication to a end-entity device when it
   should start attempting to renew its certificates.

   Prior art is that client decides, with a typical recommmendation to
   start when the remaining lifetime of the certificate is at the 50%
   point.  As typical certificate lifetimes are reduced from years to
   fractions of a year, the 50% may be far too early, and this document
   provides a way to give alternate advice.

full_text:
LAMPS Working Group                                       R. Shekh-Yusef
Internet-Draft                                                     Ciena
Intended status: Standards Track                           M. Richardson
Expires: 16 August 2026                         Sandelman Software Works
                                                            M. Ounsworth
                                                         Entrust Limited
                                                        12 February 2026


Certificate Renewal Recommendations for Enrollment over Secure Transport
                  draft-ietf-lamps-est-renewal-info-00

Abstract

   This document describes an extension to RFC7030, Enrollment over
   Secure Transport to give an indication to a end-entity device when it
   should start attempting to renew its certificates.

   Prior art is that client decides, with a typical recommmendation to
   start when the remaining lifetime of the certificate is at the 50%
   point.  As typical certificate lifetimes are reduced from years to
   fractions of a year, the 50% may be far too early, and this document
   provides a way to give alternate advice.

About This Document

   This note is to be removed before publishing as an RFC.

   Status information for this document may be found at
   https://datatracker.ietf.org/doc/draft-ietf-lamps-est-renewal-info/.

   Discussion of this document takes place on the lamps Working Group
   mailing list (mailto:spasm@ietf.org), which is archived at
   https://mailarchive.ietf.org/arch/browse/spasm/.  Subscribe at
   https://www.ietf.org/mailman/listinfo/spasm/.

   Source for this draft and an issue tracker can be found at
   https://github.com/lamps-wg/lamps-est-renewal-info.

Status of This Memo

   This Internet-Draft is submitted in full conformance with the
   provisions of BCP 78 and BCP 79.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF).  Note that other groups may also distribute
   working documents as Internet-Drafts.  The list of current Internet-
   Drafts is at https://datatracker.ietf.org/drafts/current/.



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   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet-Drafts as reference
   material or to cite them other than as "work in progress."

   This Internet-Draft will expire on 16 August 2026.

Copyright Notice

   Copyright (c) 2026 IETF Trust and the persons identified as the
   document authors.  All rights reserved.

   This document is subject to BCP 78 and the IETF Trust's Legal
   Provisions Relating to IETF Documents (https://trustee.ietf.org/
   license-info) in effect on the date of publication of this document.
   Please review these documents carefully, as they describe your rights
   and restrictions with respect to this document.  Code Components
   extracted from this document must include Revised BSD License text as
   described in Section 4.e of the Trust Legal Provisions and are
   provided without warranty as described in the Revised BSD License.

Table of Contents

   1.  Introduction  . . . . . . . . . . . . . . . . . . . . . . . .   2
   2.  Terminology . . . . . . . . . . . . . . . . . . . . . . . . .   3
   3.  Protocol Details  . . . . . . . . . . . . . . . . . . . . . .   3
     3.1.  Renewal Information Request . . . . . . . . . . . . . . .   3
     3.2.  Renewal Information Response  . . . . . . . . . . . . . .   4
       3.2.1.  Base64 Not Used . . . . . . . . . . . . . . . . . . .   4
   4.  Renewal Operations  . . . . . . . . . . . . . . . . . . . . .   4
     4.1.  Fetching Schedule . . . . . . . . . . . . . . . . . . . .   5
   5.  Privacy Considerations  . . . . . . . . . . . . . . . . . . .   5
   6.  Security Considerations . . . . . . . . . . . . . . . . . . .   6
   7.  IANA Considerations . . . . . . . . . . . . . . . . . . . . .   6
   8.  Acknowledgements  . . . . . . . . . . . . . . . . . . . . . .   6
   9.  Changelog . . . . . . . . . . . . . . . . . . . . . . . . . .   6
   10. References  . . . . . . . . . . . . . . . . . . . . . . . . .   6
     10.1.  Normative References . . . . . . . . . . . . . . . . . .   6
     10.2.  Informative References . . . . . . . . . . . . . . . . .   7
   Authors' Addresses  . . . . . . . . . . . . . . . . . . . . . . .   7

1.  Introduction

   [RFC9773], Section 1 explains why certificate lifetimes and renewal
   times need more deterministic control in the ACME [RFC8555]
   ecosystem.  Similar arguments apply to the [RFC7030] ecosystem.





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   This document extends [RFC7030] to add support for renewal
   information, by adding a new entry to the HTTP URIs for Control list
   defined in [RFC7030], Section 3.2.2

   This mechanism enables EST servers to provide suggested detailed
   renewal operations to EST clients.

   The /renewal-info URI is added, as an OPTIONAL operation, to the
   table in figure 5 in section 3.2.2 of [RFC7030].

2.  Terminology

   The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT",
   "SHOULD", "SHOULD NOT", "RECOMMENDED", "NOT RECOMMENDED", "MAY", and
   "OPTIONAL" in this document are to be interpreted as described in
   BCP 14 [RFC2119] [RFC8174] when, and only when, they appear in all
   capitals, as shown here.

3.  Protocol Details

3.1.  Renewal Information Request

   To retrieve the renewal information, the EST client uses the
   following HTTP request-line:

   GET /.well-known/est/renewal-info/<certificate-id>;

   The request includes a unique identifier for the certificate in
   question.  The unique identifier is constructed by concatenating the
   base64url encoding [RFC4648] of the keyIdentifier field of the
   certificate's Authority Key Identifier (AKI) [RFC5280] extension, the
   period character ".", and the base64url encoding of the DER-encoded
   Serial Number field (without the tag and length bytes).  All trailing
   "=" characters MUST be stripped from both parts of the unique
   identifier.

   Thus, the full request URL is constructed as follows (split onto
   multiple lines for readability), where the "||" operator indicates
   string concatenation:

       url = /.well-known/est/renewal-info
         || '/'
         || base64url(AKI keyIdentifier)
         || '.'
         || base64url(Serial)






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3.2.  Renewal Information Response

   The structure the EST RenewalInfo object is as follows:

   suggestedWindow (object, required): A JSON object with two keys,
   "start" and "end", whose values are timestamps, encoded in the format
   specified in [RFC3339], which bound the window of time in which the
   CA recommends renewing the certificate.

   For example:

       HTTP/1.1 200 OK
       Content-Type: application/json
       Retry-After: 21600

       {
         "suggestedWindow": {
           "start": "2025-01-02T04:00:00Z",
           "end": "2025-01-03T04:00:00Z"
         }
       }

3.2.1.  Base64 Not Used

   [RFC7030] mistakenly declared that all content would be base64
   encoded.  [RFC8951] clarifies that the content is to be base64
   encoded, whether or not there is a Content-Transfer-Encoding header
   present.  It further clarifies that future extensions (such as this
   document) will not use base64 encoding.  The response detailed above
   is not base64 encoded.

4.  Renewal Operations

   Clients MUST attempt renewal at a time of their choosing based on the
   suggested renewal window, obtained in the previous step.

   The following algorithm is RECOMMENDED for choosing a renewal time:

   1.  Select a uniform random time within the suggested window.

   2.  If the selected time is in the past, attempt renewal immediately.

   3.  Otherwise, if the client can schedule itself to attempt renewal
       at exactly the selected time, do so.

   4.  Otherwise, if the selected time is before the next time that the
       client would wake up normally, attempt renewal immediately.




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   5.  Otherwise, sleep until the time indicated by the Retry-After
       header and return to Step 1.

   In all cases, renewal attempts are subject to the client's existing
   error backoff and retry intervals.

   A RenewalInfo object in which the end timestamp equals or precedes
   the start timestamp is invalid.  Servers MUST NOT serve such a
   response, and clients MUST treat one as though they failed to receive
   any response from the server (e.g., retry at an appropriate interval,
   renew on a fallback schedule, etc.).

4.1.  Fetching Schedule

   The advice in [RFC9773], Section 4.3 applies:

   Clients SHOULD fetch a certificate's RenewalInfo immediately after
   issuance.

   During the lifetime of a certificate, the renewal information needs
   to be fetched frequently enough that clients learn about changes in
   the suggested window quickly, but without overwhelming the server.
   This protocol uses the Retry-After header [RFC9110] to indicate to
   clients how often to retry.  Note that in other HTTP applications,
   Retry-After often indicates the minimum time to wait before retrying
   a request.  In this protocol, it indicates the desired (i.e., both
   requested minimum and maximum) amount of time to wait.

   Clients MUST NOT check a certificate's RenewalInfo after the
   certificate has expired.  Clients MUST NOT check a certificate's
   RenewalInfo after they consider the certificate to be replaced (for
   instance, after a new certificate for the same identifiers has been
   received and configured).

5.  Privacy Considerations

   A very short certificate lifetime renewal time will cause clients to
   communicate with the EST Registrar more frequently.

   EST connections for renewals typically make use of mutually
   authenticated TLS.  When the client certificate being an IDevID, or
   the last issued certificate, often an LDevID, there is potential to
   disclose identities during this connection when using TLS 1.2.

   TLS 1.3 does not suffer from this problem, and it's use is
   RECOMMENDED as per [I-D.ietf-uta-require-tls13]





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6.  Security Considerations

   Not sure what yet.

7.  IANA Considerations

   Might need a header allocation

8.  Acknowledgements

   Many bits of text are taken from [RFC9773].

9.  Changelog

10.  References

10.1.  Normative References

   [RFC2119]  Bradner, S., "Key words for use in RFCs to Indicate
              Requirement Levels", BCP 14, RFC 2119,
              DOI 10.17487/RFC2119, March 1997,
              <https://www.rfc-editor.org/info/rfc2119>.

   [RFC3339]  Klyne, G. and C. Newman, "Date and Time on the Internet:
              Timestamps", RFC 3339, DOI 10.17487/RFC3339, July 2002,
              <https://www.rfc-editor.org/info/rfc3339>.

   [RFC4648]  Josefsson, S., "The Base16, Base32, and Base64 Data
              Encodings", RFC 4648, DOI 10.17487/RFC4648, October 2006,
              <https://www.rfc-editor.org/info/rfc4648>.

   [RFC5280]  Cooper, D., Santesson, S., Farrell, S., Boeyen, S.,
              Housley, R., and W. Polk, "Internet X.509 Public Key
              Infrastructure Certificate and Certificate Revocation List
              (CRL) Profile", RFC 5280, DOI 10.17487/RFC5280, May 2008,
              <https://www.rfc-editor.org/info/rfc5280>.

   [RFC7030]  Pritikin, M., Ed., Yee, P., Ed., and D. Harkins, Ed.,
              "Enrollment over Secure Transport", RFC 7030,
              DOI 10.17487/RFC7030, October 2013,
              <https://www.rfc-editor.org/info/rfc7030>.

   [RFC8174]  Leiba, B., "Ambiguity of Uppercase vs Lowercase in RFC
              2119 Key Words", BCP 14, RFC 8174, DOI 10.17487/RFC8174,
              May 2017, <https://www.rfc-editor.org/info/rfc8174>.






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   [RFC8951]  Richardson, M., Werner, T., and W. Pan, "Clarification of
              Enrollment over Secure Transport (EST): Transfer Encodings
              and ASN.1", RFC 8951, DOI 10.17487/RFC8951, November 2020,
              <https://www.rfc-editor.org/info/rfc8951>.

10.2.  Informative References

   [I-D.ietf-uta-require-tls13]
              Salz, R. and N. Aviram, "New Protocols Using TLS Must
              Require TLS 1.3", Work in Progress, Internet-Draft, draft-
              ietf-uta-require-tls13-12, 14 April 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-uta-
              require-tls13-12>.

   [RFC8555]  Barnes, R., Hoffman-Andrews, J., McCarney, D., and J.
              Kasten, "Automatic Certificate Management Environment
              (ACME)", RFC 8555, DOI 10.17487/RFC8555, March 2019,
              <https://www.rfc-editor.org/info/rfc8555>.

   [RFC9773]  Gable, A., "ACME Renewal Information (ARI) Extension",
              RFC 9773, DOI 10.17487/RFC9773, June 2025,
              <https://www.rfc-editor.org/info/rfc9773>.

Authors' Addresses

   Rifaat Shekh-Yusef
   Ciena
   Email: rifaat.s.ietf@gmail.com


   Michael Richardson
   Sandelman Software Works
   Email: mcr+ietf@sandelman.ca


   Mike Ounsworth
   Entrust Limited
   Email: mike.ounsworth@entrust.com













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name: draft-ietf-tls-extended-key-update
rev: 09
title: Extended Key Update for Transport Layer Security (TLS) 1.3
time: 2026-02-19T05:02:41Z
group_uri: /api/v1/group/group/1326/
expires: 2026-08-23T05:02:41Z
states: ["/api/v1/doc/state/1/", "/api/v1/doc/state/150/", "/api/v1/doc/state/38/"]
categories: ["Data formats/interop"]
tags: []
fetched_at: 2026-02-27T22:56:09.014967+00:00
abstract:
TLS 1.3 ensures forward secrecy by performing an ephemeral Diffie-
   Hellman key exchange during the initial handshake, protecting past
   communications even if a party's long-term keys (typically a private
   key with a corresponding certificate) are later compromised.  While
   the built-in KeyUpdate mechanism allows application traffic keys to
   be refreshed during a session, it does not incorporate fresh entropy
   from a new key exchange and therefore does not provide post-
   compromise security.  This limitation can pose a security risk in
   long-lived sessions, such as those found in industrial IoT or
   telecommunications environments.

   To address this, this specification defines an extended key update
   mechanism that performs a fresh Diffie-Hellman exchange within an
   active session, thereby ensuring post-compromise security.  By
   forcing attackers to exfiltrate new key material repeatedly, this
   approach mitigates the risks associated with static key compromise.
   Regular renewal of session keys helps contain the impact of such
   compromises.  The extension is applicable to both TLS 1.3 and DTLS
   1.3.

full_text:
Transport Layer Security                                   H. Tschofenig
Internet-Draft                                                   Siemens
Updates: 9261, 8446 (if approved)                               M. Tüxen
Intended status: Standards Track       Münster Univ. of Applied Sciences
Expires: 23 August 2026                                         T. Reddy
                                                                   Nokia
                                                                S. Fries
                                                                 Siemens
                                                            Y. Rosomakho
                                                                 Zscaler
                                                        19 February 2026


       Extended Key Update for Transport Layer Security (TLS) 1.3
                 draft-ietf-tls-extended-key-update-09

Abstract

   TLS 1.3 ensures forward secrecy by performing an ephemeral Diffie-
   Hellman key exchange during the initial handshake, protecting past
   communications even if a party's long-term keys (typically a private
   key with a corresponding certificate) are later compromised.  While
   the built-in KeyUpdate mechanism allows application traffic keys to
   be refreshed during a session, it does not incorporate fresh entropy
   from a new key exchange and therefore does not provide post-
   compromise security.  This limitation can pose a security risk in
   long-lived sessions, such as those found in industrial IoT or
   telecommunications environments.

   To address this, this specification defines an extended key update
   mechanism that performs a fresh Diffie-Hellman exchange within an
   active session, thereby ensuring post-compromise security.  By
   forcing attackers to exfiltrate new key material repeatedly, this
   approach mitigates the risks associated with static key compromise.
   Regular renewal of session keys helps contain the impact of such
   compromises.  The extension is applicable to both TLS 1.3 and DTLS
   1.3.

Status of This Memo

   This Internet-Draft is submitted in full conformance with the
   provisions of BCP 78 and BCP 79.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF).  Note that other groups may also distribute
   working documents as Internet-Drafts.  The list of current Internet-
   Drafts is at https://datatracker.ietf.org/drafts/current/.




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   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet-Drafts as reference
   material or to cite them other than as "work in progress."

   This Internet-Draft will expire on 23 August 2026.

Copyright Notice

   Copyright (c) 2026 IETF Trust and the persons identified as the
   document authors.  All rights reserved.

   This document is subject to BCP 78 and the IETF Trust's Legal
   Provisions Relating to IETF Documents (https://trustee.ietf.org/
   license-info) in effect on the date of publication of this document.
   Please review these documents carefully, as they describe your rights
   and restrictions with respect to this document.  Code Components
   extracted from this document must include Revised BSD License text as
   described in Section 4.e of the Trust Legal Provisions and are
   provided without warranty as described in the Revised BSD License.

Table of Contents

   1.  Introduction  . . . . . . . . . . . . . . . . . . . . . . . .   3
   2.  Terminology and Requirements Language . . . . . . . . . . . .   4
   3.  Negotiating the Extended Key Update . . . . . . . . . . . . .   5
   4.  Extended Key Update Messages  . . . . . . . . . . . . . . . .   5
   5.  TLS 1.3 Considerations  . . . . . . . . . . . . . . . . . . .   8
     5.1.  TLS 1.3 Extended Key Update Example . . . . . . . . . . .  10
   6.  DTLS 1.3 Considerations . . . . . . . . . . . . . . . . . . .  11
     6.1.  DTLS 1.3 Extended Key Update Example  . . . . . . . . . .  14
   7.  Updating Traffic Keys . . . . . . . . . . . . . . . . . . . .  16
   8.  Post-Quantum Cryptography Considerations  . . . . . . . . . .  18
   9.  SSLKEYLOGFILE Update  . . . . . . . . . . . . . . . . . . . .  19
   10. Exporter  . . . . . . . . . . . . . . . . . . . . . . . . . .  19
     10.1.  Post-Compromise Security for the Initial Exporter
            Secret . . . . . . . . . . . . . . . . . . . . . . . . .  20
     10.2.  Exporter Usage After Extended Key Update . . . . . . . .  21
   11. Use of Post-Handshake Authentication and Exported
           Authenticators with Extended Key Update . . . . . . . . .  22
     11.1.  Post-Handshake Certificate-Based Client
            Authentication . . . . . . . . . . . . . . . . . . . . .  22
     11.2.  Exported Authenticators  . . . . . . . . . . . . . . . .  23
   12. Security Considerations . . . . . . . . . . . . . . . . . . .  23
     12.1.  Scope of Key Compromise  . . . . . . . . . . . . . . . .  24
     12.2.  Post-Compromise Security . . . . . . . . . . . . . . . .  24
     12.3.  Denial-of-Service (DoS)  . . . . . . . . . . . . . . . .  25
     12.4.  Operational Guidance . . . . . . . . . . . . . . . . . .  25



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   13. IANA Considerations . . . . . . . . . . . . . . . . . . . . .  25
     13.1.  TLS Flags  . . . . . . . . . . . . . . . . . . . . . . .  25
     13.2.  TLS HandshakeType  . . . . . . . . . . . . . . . . . . .  25
     13.3.  ExtendedKeyUpdate Message Subtypes Registry  . . . . . .  26
     13.4.  SSLKEYLOGFILE labels . . . . . . . . . . . . . . . . . .  26
   14. References  . . . . . . . . . . . . . . . . . . . . . . . . .  27
     14.1.  Normative References . . . . . . . . . . . . . . . . . .  27
     14.2.  Informative References . . . . . . . . . . . . . . . . .  28
   Appendix A.  Acknowledgments  . . . . . . . . . . . . . . . . . .  30
   Appendix B.  State Machines . . . . . . . . . . . . . . . . . . .  30
     B.1.  TLS 1.3 State Machines  . . . . . . . . . . . . . . . . .  31
       B.1.1.  Initiator State Machine . . . . . . . . . . . . . . .  31
       B.1.2.  Responder State Machine . . . . . . . . . . . . . . .  32
     B.2.  DTLS 1.3 State Machines . . . . . . . . . . . . . . . . .  33
       B.2.1.  Terms and Abbreviations . . . . . . . . . . . . . . .  34
       B.2.2.  State Machine (Initiator) . . . . . . . . . . . . . .  34
       B.2.3.  State Machine (Responder) . . . . . . . . . . . . . .  36
   Appendix C.  Overview of Security Goals . . . . . . . . . . . . .  38
     C.1.  Post-Compromise Security (PCS)  . . . . . . . . . . . . .  39
     C.2.  Key Freshness and Cryptographic Independence  . . . . . .  39
     C.3.  Elimination of Standard KeyUpdate . . . . . . . . . . . .  39
     C.4.  Detecting Divergent Key State . . . . . . . . . . . . . .  39
   Authors' Addresses  . . . . . . . . . . . . . . . . . . . . . . .  40

1.  Introduction

   The Transport Layer Security (TLS) 1.3 protocol provides forward
   secrecy by using an ephemeral Diffie-Hellman (DHE) key exchange
   during the initial handshake.  This ensures that encrypted
   communication remains confidential even if an attacker later obtains
   a party's long-term private key, protecting against passive
   adversaries who record encrypted traffic for later decryption.

   TLS 1.3 also includes a KeyUpdate mechanism that allows traffic keys
   to be refreshed during an established session.  However, this
   mechanism does not provide post-compromise security, as it applies
   only a key derivation function to the previous application traffic
   key as input.  While this design is generally sufficient for short-
   lived connections, it may present security limitations in scenarios
   where sessions persist for extended periods, such as in industrial
   IoT or telecommunications systems, where continuous availability is
   critical and session renegotiation or resumption is impractical.

   Earlier versions of TLS supported session renegotiation, which
   allowed peers to negotiate fresh keying material, including
   performing new Diffie-Hellman exchanges during the session lifetime.
   Due to protocol complexity and known vulnerabilities, renegotiation
   was first restricted by [TLS-RENEGOTIATION] and ultimately removed in



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   TLS 1.3.  While the KeyUpdate message was introduced to offer limited
   rekeying functionality, it does not fulfill the same cryptographic
   role as renegotiation and cannot refresh the TLS main secret and
   consequently cannot derive new secrets from fresh DHE input.

   Security guidance from national agencies, such as ANSSI (France
   [ANSSI]), recommends the periodic renewal of cryptographic keys
   during long-lived sessions to limit the impact of key compromise.
   This approach encourages designs that force an attacker to perform
   dynamic key exfiltration, as defined in [CONFIDENTIALITY].  Dynamic
   key exfiltration refers to attack scenarios where an adversary must
   repeatedly extract fresh keying material to maintain access to
   protected data, increasing operational cost and risk for the
   attacker.  In contrast, static key exfiltration, where the TLS main
   secret is extracted once and reused, poses a greater long-term
   threat, especially when session keys are not refreshed with fresh key
   exchange input rather than key derivation.

   This specification defines a TLS extension that introduces an
   extended key update mechanism.  Unlike the standard key update, this
   mechanism allows peers to perform a fresh Diffie-Hellman exchange
   within an active session using one of the groups negotiated during
   the initial handshake.  By periodically rerunning (EC)DHE, this
   extension enables the derivation of new traffic keys that are
   independent of prior key material.  As noted in Appendix F of [TLS],
   this approach mitigates the risk of static key exfiltration and
   shifts the attacker burden toward dynamic key exfiltration.

   The proposed extension is applicable to both TLS 1.3 [TLS] and DTLS
   1.3 [DTLS].  For clarity, the term "TLS" is used throughout this
   document to refer to both protocols unless otherwise specified.

2.  Terminology and Requirements Language

   The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT",
   "SHOULD", "SHOULD NOT", "RECOMMENDED", "NOT RECOMMENDED", "MAY", and
   "OPTIONAL" in this document are to be interpreted as described in
   BCP 14 [RFC2119] [RFC8174] when, and only when, they appear in all
   capitals, as shown here.

   To distinguish the key update procedure defined in [TLS] from the key
   update procedure specified in this document, we use the terms
   "standard key update" and "extended key update", respectively.

   In this document, we use the term post-compromise security, as
   defined in [CCG16].  We assume that an adversary may obtain access to
   the application traffic keys.




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   Unless otherwise specified, all references to traffic keys in this
   document refer to application traffic keys and because the Extended
   Key Update procedure occurs after the handshake phase has completed,
   no handshake traffic keys are involved.

   In this document, send key refers to the [sender]_write_key, and
   receive key refers to the [receiver]_write_key.  These keys are
   derived from the active client_application_traffic_secret_N and
   server_application_traffic_secret_N, as defined in (D)TLS 1.3 [TLS],
   and are replaced with new ones after each successful Extended Key
   Update.

3.  Negotiating the Extended Key Update

   Client and servers use the TLS flags extension [TLS-FLAGS] to
   indicate support for the functionality defined in this document.  We
   call this flag "Extended_Key_Update" flag.

   The "Extended_Key_Update" flag proposed by the client in the
   ClientHello (CH) MUST be acknowledged in the EncryptedExtensions
   (EE), if the server also supports the functionality defined in this
   document and is configured to use it.

   If the "Extended_Key_Update" flag is not set, servers ignore any of
   the functionality specified in this document and applications that
   require post-compromise security will have to initiate a full
   handshake.

4.  Extended Key Update Messages

   If the client and server agree to use the extended key update
   mechanism, the standard key update MUST NOT be used.  In this case,
   the extended key update fully replaces the standard key update
   functionality.

   Implementations that receive a classic KeyUpdate message after
   successfully negotiating the Extended Key Update functionality MUST
   terminate the connection with an "unexpected_message" alert.

   The extended key update messages are signaled in a new handshake
   message named ExtendedKeyUpdate (EKU), with an internal uint8 message
   subtype indicating its role.  This specification defines three
   ExtendedKeyUpdate message subtypes:

   *  key_update_request (0)

   *  key_update_response (1)




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   *  new_key_update (2)

   New ExtendedKeyUpdate message subtypes are assigned by IANA as
   described in Section 13.3.

   A TLS peer which receives a ExtendedKeyUpdate with an unexpected
   message subtype MUST abort the connection with an
   "unexpected_message" alert.

   The extended key update process can be initiated by either peer after
   it has sent a Finished message.  Implementations that receive an
   ExtendedKeyUpdate message prior to the sender having sent Finished
   MUST terminate the connection with an "unexpected_message" alert.

   The KeyShareEntry carried in a key_update_request and in a
   key_update_response MUST use the group that was negotiated by the
   client and server during the initial handshake.  An implementation
   that receives an algorithm other than previously negotiated MUST
   terminate the connection with an "illegal_parameter" alert.

   Figure 1 shows the interaction graphically.  First, support for the
   functionality in this specification is negotiated in the ClientHello
   and the EncryptedExtensions messages.  Then, the ExtendedKeyUpdate
   exchange is sent to update the application traffic secrets.

   The extended key update exchange is performed between the initiator
   and the responder; either the TLS client or the TLS server may act as
   initiator.























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        Client                                           Server

 Key  ^ ClientHello
 Exch | + key_share
      | + signature_algorithms
      v + extended_key_update   -------->
                                                   ServerHello  ^ Key
                                                   + key_share  | Exch
                                                                v
                                         {EncryptedExtensions   ^ Server
                                        + extended_key_update}  | Params
                                          {CertificateRequest}  v
                                                 {Certificate}  ^
                                           {CertificateVerify}  | Auth
                                                    {Finished}  v
                                <--------
      ^ {Certificate}
 Auth | {CertificateVerify}
      v {Finished}              -------->
        [Application Data]N     <------->  [Application Data]N
                                   ...
   [EKU(key_update_request)]N   -------->
                                <--------  [EKU(key_update_response)]N
       [EKU(new_key_update)]N   -------->
                                   ...
        [Application Data]N+1   <------->  [Application Data]N+1

 Legend:

     +   Indicates noteworthy extensions sent in the
     previously noted message.

     -   Indicates optional or situation-dependent
     messages/extensions that are not always sent.

     () Indicates messages protected using keys
     derived from a client_early_traffic_secret.

     {} Indicates messages protected using keys
     derived from a [sender]_handshake_traffic_secret.

     []N Indicates messages protected using keys
     derived from [sender]_application_traffic_secret_N.

       Figure 1: Extended Key Update Message Exchange in TLS 1.3.

   The ExtendedKeyUpdate wire format is:




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   enum {
      key_update_request(0),
      key_update_response(1),
      new_key_update(2),
      (255)
   } ExtendedKeyUpdateType;

   struct {
      ExtendedKeyUpdateType eku_type;
      select (eku_type) {
         case key_update_request: {
             KeyShareEntry key_share;
         }
         case key_update_response: {
             KeyShareEntry key_share;
         }
         case new_key_update: {
             /* empty */
         }
      };
   } ExtendedKeyUpdate;

   Fields:

   *  eku_type: the subtype of the ExtendedKeyUpdate message.

   *  key_share: key share information.  The contents of this field is
      determined by the specified group and its corresponding definition
      (see Section 4.2.8 of [TLS]).

5.  TLS 1.3 Considerations

   The following steps are taken by a TLS 1.3 implementation; the steps
   executed with DTLS 1.3 differ slightly.

   1.  The initiator sends ExtendedKeyUpdate(key_update_request)
       carrying a KeyShareEntry.  While an extended key update is in
       progress, the initiator MUST NOT initiate another key update.

   2.  Upon receipt, the responder sends its own KeyShareEntry in a
       ExtendedKeyUpdate(key_update_response) message.  While an
       extended key update is in progress, the responder MUST NOT
       initiate another key update.  The responder MAY defer sending a
       response if system load or resource constraints prevent immediate
       processing.  In such cases, the response MUST be sent once
       sufficient resources become available.





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   3.  After the responder sends ExtendedKeyUpdate(key_update_response)
       it MUST update its send keys.

   4.  Upon receipt of an ExtendedKeyUpdate(key_update_response) the
       initiator derives the new secrets from the exchanged key shares.
       The initiator then updates its receive keys and sends an empty
       ExtendedKeyUpdate(new_key_update) message to complete the
       process.  The initiator MUST NOT defer derivation of the secrets
       and sending the ExtendedKeyUpdate(new_key_update) message as it
       would stall the communication.

   5.  After sending ExtendedKeyUpdate(new_key_update) initiator MUST
       update its send keys.

   6.  After receiving the initiator’s
       ExtendedKeyUpdate(new_key_update), the responder MUST update its
       receive keys.

   Both initiator and responder MUST encrypt their ExtendedKeyUpdate
   messages with the old keys.  Responder MUST ensure that the
   ExtendedKeyUpdate(new_key_update) encrypted with the old key is
   received before accepting any messages encrypted with the new key.

   If TLS peers independently initiate the extended key update and the
   requests cross in flight, the ExtendedKeyUpdate(key_update_request)
   with the lower lexicographic order of the key_exchange value in
   KeyShareEntry MUST be ignored.  This prevents each side from
   advancing keys by two generations.  If the tie-break comparison
   yields equality (an event that should be impossible for genuine
   asymmetric key pairs), the endpoint MUST treat this as a protocol
   violation, send an "unexpected_message" alert, and close the
   connection.

   The handshake framing uses a single HandshakeType for this message
   (see Figure 2).
















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         struct {
             HandshakeType msg_type;    /* handshake type */
             uint24 length;             /* bytes in message */
             select (Handshake.msg_type) {
                 case client_hello:          ClientHello;
                 case server_hello:          ServerHello;
                 case end_of_early_data:     EndOfEarlyData;
                 case encrypted_extensions:  EncryptedExtensions;
                 case certificate_request:   CertificateRequest;
                 case certificate:           Certificate;
                 case certificate_verify:    CertificateVerify;
                 case finished:              Finished;
                 case new_session_ticket:    NewSessionTicket;
                 case key_update:            KeyUpdate;
                 case extended_key_update:   ExtendedKeyUpdate;
             };
         } Handshake;

                   Figure 2: TLS 1.3 Handshake Structure.

5.1.  TLS 1.3 Extended Key Update Example

   While Figure 1 shows the high-level interaction between a TLS 1.3
   client and server, this section shows an example message exchange
   with information about the updated keys added.

   There are two phases:

   1.  The support for the functionality in this specification is
       negotiated in the ClientHello and the EncryptedExtensions
       messages.

   2.  Once the initial handshake is completed, a key update can be
       triggered.

   Figure 3 provides an overview of the exchange starting with the
   initial negotiation followed by the key update.














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        Client                                           Server

 Key  ^ ClientHello
 Exch | + key_share
      | + signature_algorithms
      v + extended_key_update
                              -------->
                                                   ServerHello  ^ Key
                                                   + key_share  | Exch
                                                                v
                                         {EncryptedExtensions   ^ Server
                                        + extended_key_update}  | Params
                                          {CertificateRequest}  v
                                                 {Certificate}  ^
                                           {CertificateVerify}  | Auth
                                                    {Finished}  v
                                <--------
      ^ {Certificate}
 Auth | {CertificateVerify}
      v {Finished}              -------->
                                   ...
                               some time later
                                   ...
  [EKU(key_update_request       -------->
        (with key_share))]
                                <-------- [EKU(key_update_response
                                            (with key_share))]
                                         # Server derives new secrets
                                         # and updates SEND keys here
 # Client derives new secrets
 # and updates RECEIVE keys here
         [EKU(new_key_update)]  -------->
 # Client updates SEND keys here
                                     # Server updates RECEIVE keys here

                 Figure 3: Extended Key Update Example.

6.  DTLS 1.3 Considerations

   Unlike TLS 1.3, DTLS 1.3 implementations must take into account that
   handshake messages are not transmitted over a reliable transport
   protocol.









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   EKU messages MUST be transmitted reliably, like other DTLS handshake
   messages.  If necessary, EKU messages MAY be fragmented as described
   in Section 5.5 of [DTLS].  Due to the possibility of an
   ExtendedKeyUpdate messages being lost and thereby preventing the
   sender of that message from updating its keying material, receivers
   MUST retain the pre-update keying material until receipt and
   successful decryption of a message using the new keys.

   Due to packet loss and/or reordering, DTLS 1.3 peers MAY receive
   records from an earlier epoch.  If the necessary keys are available,
   implementations SHOULD attempt to process such records; however, they
   MAY choose to discard them.  The exchange has the following steps:

   1.  The initiator sends an ExtendedKeyUpdate(key_update_request)
       message, which contains a key share.  While an extended key
       update is in progress, the initiator MUST NOT initiate further
       key updates.  This message is subject to DTLS handshake
       retransmission, but delivery is only confirmed when the initiator
       either receives the corresponding
       ExtendedKeyUpdate(key_update_response) or an ACK.

   2.  Upon receipt, the responder sends its own KeyShareEntry in a
       ExtendedKeyUpdate(key_update_response) message.  While an
       extended key update is in progress, the responder MUST NOT
       initiate further key updates.  The responder MAY defer sending a
       response if system load or resource constraints prevent immediate
       processing.  In such cases, the responder MUST acknowledge
       receipt of the key_update_request with an ACK and, once
       sufficient resources become available, retransmit the
       key_update_response until it is acknowledged by the initiator.

   3.  On receipt of ExtendedKeyUpdate(key_update_response) the
       initiator derives a secret key based on the exchanged key shares.
       This message also serves as an implicit acknowledgment of the
       initiator's ExtendedKeyUpdate(key_update_request), so no separate
       ACK is required.  The initiator MUST update its receive keys and
       epoch value.  The initiator MUST NOT defer derivation of the
       secrets.

   4.  The initiator transmits an ExtendedKeyUpdate(new_key_update)
       message.  This message is subject to DTLS retransmission until
       acknowledged.

   5.  The responder MUST acknowledge the received message by sending an
       ACK message.






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   6.  After the responder receives the initiator's
       ExtendedKeyUpdate(new_key_update), the responder MUST update its
       send key and epoch value.  With the receipt of that message, the
       responder MUST also update its receive keys.

   7.  On receipt of the ACK message, the initiator updates its send key
       and epoch value.  If this ACK is not received, the initiator re-
       transmits its ExtendedKeyUpdate(new_key_update) until ACK is
       received.  The key update is complete once this ACK is processed
       by the initiator.

   The handshake framing uses a single HandshakeType for this message
   (see Figure 4).






































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          enum {
              client_hello(1),
              server_hello(2),
              new_session_ticket(4),
              end_of_early_data(5),
              encrypted_extensions(8),
              request_connection_id(9),
              new_connection_id(10),
              certificate(11),
              certificate_request(13),
              certificate_verify(15),
              finished(20),
              key_update(24),
              extended_key_update(TBD),  /* new */
              message_hash(254),
              (255)
          } HandshakeType;

          struct {
              HandshakeType msg_type;    /* handshake type */
              uint24 length;             /* bytes in message */
              uint16 message_seq;        /* DTLS-required field */
              uint24 fragment_offset;    /* DTLS-required field */
              uint24 fragment_length;    /* DTLS-required field */
              select (msg_type) {
                  case client_hello:          ClientHello;
                  case server_hello:          ServerHello;
                  case end_of_early_data:     EndOfEarlyData;
                  case encrypted_extensions:  EncryptedExtensions;
                  case certificate_request:   CertificateRequest;
                  case certificate:           Certificate;
                  case certificate_verify:    CertificateVerify;
                  case finished:              Finished;
                  case new_session_ticket:    NewSessionTicket;
                  case key_update:            KeyUpdate;
                  case extended_key_update:   ExtendedKeyUpdate;
                  case request_connection_id: RequestConnectionId;
                  case new_connection_id:     NewConnectionId;
              } body;
          } DTLSHandshake;

                  Figure 4: DTLS 1.3 Handshake Structure.

6.1.  DTLS 1.3 Extended Key Update Example

   The following example illustrates a successful extended key update,
   including how the epochs change during the exchange.




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   Client                            Server
   (Initiator)                       (Responder)

     /---------------------------------------\
    |           Initial Handshake             |
     \---------------------------------------/

   [C: tx=3, rx=3]                   [S: tx=3, rx=3]
   [Application Data]               -------->
   [C: tx=3, rx=3]                   [S: tx=3, rx=3]

   [C: tx=3, rx=3]                   [S: tx=3, rx=3]
                                     <--------   [Application Data]
   [C: tx=3, rx=3]                   [S: tx=3, rx=3]

     /---------------------------------------\
    |           Some time later ...           |
     \---------------------------------------/

   [C: tx=3, rx=3]                   [S: tx=3, rx=3]
   [EKU(key_update_request)]  -------->
                                     # no epoch change yet

                              <-------- [EKU(key_update_response)]
   # Sent under OLD epoch. Server does NOT bump yet.

   # Step 3: initiator bumps RECEIVE epoch on
   # receiving key update response-in:
   # (rx:=rx+1; tx still old)
   [C: tx=3, rx=4]                   [S: tx=3, rx=3]

   [EKU(new_key_update)]      -------->
   # Sender's NKU is tagged with OLD tx (3).

   # Epoch switch point:
   # Step 6: responder bumps BOTH tx and rx on NKU-in:
   [C: tx=3, rx=4]                   [S: tx=4, rx=4]

                              <-------- [ACK] (tag=new)

   # Step 7: initiator bumps SEND epoch on ACK-in:
   [C: tx=4, rx=4]                   [S: tx=4, rx=4]

   [Application Data]         -------->
   [C: tx=4, rx=4]                   [S: tx=4, rx=4]

                              <--------   [Application Data]
   [C: tx=4, rx=4]                   [S: tx=4, rx=4]



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     Figure 5: Example DTLS 1.3 Extended Key Update: Message Exchange.

   Figure 6 shows the steps, the message in flight, and the epoch
   changes on both sides.  The A/B -> X/Y notation indicates the change
   of epoch values for tx/rx before and after the message transmission.

   +--------------------+----------------+--------------+
   | Message            | Client tx/rx   | Server tx/rx |
   +--------------------+----------------+--------------+
   | APP ------------>  | 3/3 -> 3/3     | 3/3 -> 3/3   |
   | <------------ APP  | 3/3 -> 3/3     | 3/3 -> 3/3   |
   | req -------------> | 3/3 -> 3/3     | 3/3 -> 3/3   |
   | <------------ resp | 3/3 -> 3/4     | 3/3 -> 3/3   | <- step 3
   | NKU  ------------> | 3/4 -> 3/4     | 3/3 -> 4/4   | <- step 6
   | <------------- ACK | 3/4 -> 4/4     | 4/4 -> 4/4   | <- step 7
   | APP -------------> | 4/4 -> 4/4     | 4/4 -> 4/4   |
   | <------------- APP | 4/4 -> 4/4     | 4/4 -> 4/4   |
   +--------------------+----------------+--------------+

       Figure 6: Example DTLS 1.3 Extended Key Update: Epoch Changes.

7.  Updating Traffic Keys

   When the extended key update message exchange is completed both peers
   have successfully updated their application traffic keys.  The key
   derivation function described in this document is used to perform
   this update.

   The design of the key derivation function for computing the next
   generation of application_traffic_secret is motivated by the desire
   to include

   *  a secret derived from the (EC)DHE exchange (or from the hybrid key
      exchange / PQ-KEM exchange),

   *  a secret that allows the new key exchange to be cryptographically
      bound to the previously established secret,

   *  a transcript hash that is updated after each Extended Key Update
      exchange by hashing together the previous transcript hash value
      and the current ExtendedKeyUpdate(key_update_request) and
      ExtendedKeyUpdate(key_update_response) messages, which contain the
      key shares, thereby binding the encapsulated shared secret
      ciphertext to the IKM in the case of PQ/T hybrid or PQC key
      exchange and cryptographically binding the newly derived secrets
      to the prior handshake transcript and all preceding EKU exchanges,
      as well as to the current EKU exchange, and




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   *  new label strings to distinguish it from the key derivation used
      in TLS 1.3.

   The transcript_hash_N denotes the transcript hash value associated
   with generation N.  During each Extended Key Update exchange, the
   transcript hash value for the next generation is computed as follows:

    transcript_hash_N+1 = Transcript-Hash(transcript_hash_N ||
                                          EKU(key_update_request) ||
                                          EKU(key_update_response))

   Once transcript_hash_N+1 has been computed, transcript_hash_N can be
   deleted.  No prior transcript hash values need to be retained for
   future EKU exchanges. transcript_hash_0 denotes the transcript hash
   of the initial TLS handshake, covering all messages from the
   ClientHello up to and including the client Finished message.

   The following diagram shows the key derivation hierarchy.

          Main Secret N
                |
                v
          Derive-Secret(., "derived", "")
                |
                v
    (EC)DHE -> HKDF-Extract = Main Secret N+1
                |
                +-----> Derive-Secret(., "c ap traffic",
                |                     transcript_hash_N+1)
                |                = client_application_traffic_secret_N+1
                |
                +-----> Derive-Secret(., "s ap traffic",
                |                     transcript_hash_N+1)
                |                = server_application_traffic_secret_N+1
                |
                +-----> Derive-Secret(., "exp master",
                |                     transcript_hash_N+1)
                |                = exporter_secret_N+1
                |
                +-----> Derive-Secret(., "res master",
                |                     transcript_hash_N+1)
                                 = resumption_main_secret_N+1

   During the initial handshake, the Main Secret is generated (see
   Section 7.1 of [TLS]).  Since the Main Secret is discarded during the
   key derivation procedure, a derived value is stored.  This stored
   value then serves as the input salt to the first key update procedure
   that incorporates the ephemeral (EC)DHE- established value as input



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   keying material (IKM) to produce main_secret_{N+1}. The derived value
   from this new master secret serves as input salt to the subsequent
   key update procedure, which also incorporates a fresh ephemeral
   (EC)DHE value as IKM.  This process is repeated for each additional
   key update procedure.

   The traffic keys are re-derived from
   client_application_traffic_secret_N+1 and
   server_application_traffic_secret_N+1, as described in Section 7.3 of
   [TLS].

   Once client_/server_application_traffic_secret_N+1 and its associated
   traffic keys have been computed, implementations SHOULD delete
   client_/server_application_traffic_secret_N and its associated
   traffic keys as soon as possible.  Note: The client_/
   server_application_traffic_secret_N and its associated traffic keys
   can only be deleted by the responder after receiving the
   ExtendedKeyUpdate(new_key_update) message.

   Once client_/server_application_traffic_secret_N+1 and the
   corresponding traffic keys are in use, all subsequent records,
   including alerts and post-handshake messages MUST be protected using
   those keys.

   When using this extension, it is important to consider its
   interaction with ticket-based session resumption.  If resumption
   occurs without a new (EC)DH exchange that provides forward secrecy,
   an attacker could potentially revert the security context to an
   earlier state, thereby negating the benefits of the extended key
   update.  To preserve the security guarantees provided by key updates,
   endpoints MUST either invalidate any session tickets issued prior to
   the key update or ensure that resumption always involves a fresh
   (EC)DH exchange.

   If session tickets cannot be stored securely, developers SHOULD
   consider disabling ticket-based resumption in their deployments.
   While this approach may impact performance, it provides improved
   security properties.

8.  Post-Quantum Cryptography Considerations

   Hybrid key exchange refers to the simultaneous use of multiple key
   exchange algorithms, with the resulting shared secret derived by
   combining the outputs of each.  The goal of this approach is to
   maintain security even if all but one of the component algorithms are
   later found to be vulnerable.





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   The transition to post-quantum cryptography has motivated the
   adoption of hybrid key exchanges in TLS, as described in
   [TLS-HYBRID].  Specific hybrid groups have been registered in
   [TLS-ECDHE-MLKEM].  When hybrid key exchange is used, the
   key_exchange field of each KeyShareEntry in the initial handshake is
   formed by concatenating the key_exchange fields of the constituent
   algorithms.  This same approach is reused during the Extended Key
   Update, when new key shares are exchanged.

   The specification in [TLS-MLKEM] registers the lattice-based ML-KEM
   algorithm and its variants, such as ML-KEM-512, ML-KEM-768 and ML-
   KEM-1024.  The KEM encapsulation key or KEM ciphertext is represented
   as a 'KeyShareEntry' field.  This same approach is reused during the
   Extended Key Update, when new key shares are exchanged.

9.  SSLKEYLOGFILE Update

   As a successful extended key update exchange invalidates previous
   secrets, SSLKEYLOGFILE [TLS-KEYLOGFILE] needs to be populated with
   new entries.  As a result, two additional secret labels are utilized
   in the SSLKEYLOGFILE:

   1.  CLIENT_TRAFFIC_SECRET_N+1: identifies the
       client_application_traffic_secret_N+1 in the key schedule

   2.  SERVER_TRAFFIC_SECRET_N+1: identifies the
       server_application_traffic_secret_N+1 in the key schedule

   3.  EXPORTER_SECRET_N+1: identifies the exporter_secret_N+1 in the
       key schedule

   Similar to other entries in the SSLKEYLOGFILE, the label is followed
   by the 32-byte value of the Random field from the ClientHello message
   that established the TLS connection, and the corresponding secret
   encoded in hexadecimal.

   SSLKEYLOGFILE entries for the extended key update MUST NOT be
   produced if SSLKEYLOGFILE was not used for other secrets in the
   handshake.

   Note that each successful Extended Key Update invalidates all
   previous SSLKEYLOGFILE secrets including past iterations of
   CLIENT_TRAFFIC_SECRET_, SERVER_TRAFFIC_SECRET_ and EXPORTER_SECRET_.

10.  Exporter






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10.1.  Post-Compromise Security for the Initial Exporter Secret

   The TLS 1.3 Key Schedule, see Figure 5 of [TLS], derives the
   exporter_secret from the main secret.  This exporter_secret is static
   for the lifetime of the connection and is not updated by a standard
   key update.

   A core design goal of this specification is not met if the
   exporter_secret does not change.  Therefore, this document defines an
   exporter interface that derives a fresh exporter secret whenever new
   application traffic keys are updated through the EKU.

   If the initial exporter secret for this new interface were identical
   to exporter_secret, then compromising exporter_secret at any point
   during the lifetime of the connection would enable an attacker to
   recompute all exporter outputs derived from it.  This would break
   post-compromise security for exported keying material.

   Therefore, the initial exporter secret used by the exporter interface
   defined in this document, i.e., the exporter output available prior
   to the first Extended Key Update, MUST be distinct from the
   exporter_secret.  This separation ensures that compromise of the TLS
   exporter interface does not compromise outputs derived from the
   exporter interface defined in this document.

   Prior to the first Extended Key Update, the exporter interface
   provides an initial exporter secret, denoted exporter_secret_0.  This
   secret is derived from the TLS main secret and the handshake
   transcript, but is cryptographically independent of the TLS
   exporter_secret.  It is computed as follows:

   exporter_secret_0 =
   Derive-Secret(Main Secret,
   "exporter eku",
   Transcript-Hash(ClientHello..server Finished))

   Applications that require post-compromise security MUST use the
   exporter interface defined in this document.  This exporter interface
   is independent of the TLS exporter defined in Section 7.5 of [TLS],
   which continues to use a static exporter_secret for the lifetime of
   the connection for compatiblity with "legacy" applications.










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10.2.  Exporter Usage After Extended Key Update

   Protocols such as DTLS-SRTP and DTLS-over-SCTP rely on TLS or DTLS
   for key establishment, but reuse portions of the derived keying
   material for their own specific purposes.  These protocols use the
   TLS exporter defined in Section 7.5 of [TLS].  Exporters are also
   used for deriving authentication related values such as nonces, as
   described in [RFC9729].

   Once the Extended Key Update mechanism is complete, such protocols
   would need to use the newly derived exporter secret to generate
   Exported Keying Material (EKM) to protect packets.  The "sk" derived
   in the Section 7 will be used as the "Secret" in the exporter
   function, defined in Section 7.5 of [TLS], to generate EKM, ensuring
   that the exported keying material is aligned with the updated
   security context.  The newly derived exporter secret is
   cryptographically independent of previous exporter secrets.

   When a new exporter secret becomes active following a successful
   Extended Key Update, the TLS or DTLS implementation would have to
   provide an asynchronous notification to the application indicating
   that:

   *  A new epoch has become active, and the (D)TLS implementation can
      include the corresponding epoch identifier.  Applications
      receiving an epoch identifier can use it to request keying
      material for that specific epoch through an epoch-aware exporter
      interface.  In TLS, this identifier represents a local logical
      counter that may differ between peers.

   Applications are notified that a new epoch is active only after both
   peers have completed the Extended Key Update exchange and switched to
   the new traffic keys.

   *  In TLS, the initiator triggers notification after Step 5 in
      Section 5, and the responder triggers notification after Step 4 in
      Section 5.

   *  In DTLS, the initiator triggers notification to the application
      after Step 7 in Section 6, and the responder triggers notification
      after Step 9 in Section 6.

   The corresponding EKM is obtained by the application through the TLS/
   DTLS exporter interface using its chosen label and context values as
   defined in Section 4 of [RFC5705].






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   To prevent desynchronization, the application will have to retain
   both the previous and the newly derived exporter secrets for a short
   period.  For TLS, the previous exporter secret would be discarded
   once data derived from the new exporter has been successfully
   processed, and no records protected with the old exporter secret are
   expected to arrive.  For DTLS, the previous exporter secret needs to
   be retained until the retention timer expires for the prior epoch, to
   allow for processing of packets that may arrive out of order.  The
   retention policy for exporter secrets is application-specific.  For
   example, in DTLS-SRTP, the application might retain the previous
   exporter secret until its replay window no longer accepts packets
   protected with keys derived from that secret, as described in
   Section 3.3.2 of [RFC3711].

11.  Use of Post-Handshake Authentication and Exported Authenticators
     with Extended Key Update

   EKU provides fresh traffic secrets, but EKU alone does not
   authenticate that both endpoints derived the same updated keys.  An
   attacker that temporarily compromises an endpoint may later act as an
   active MitM capable of interfering with the EKU exchange.  Such an
   attacker can cause the peers to transition to divergent traffic
   secrets without detection, but cannot compromise the endpoint to
   derive secrets after the new epoch is established.  To confirm that
   both peers transitioned to the same new key state, TLS 1.3 provides
   two mechanisms: Post-Handshake Certificate-Based Client
   Authentication and Exported Authenticators [RFC9261].

11.1.  Post-Handshake Certificate-Based Client Authentication

   When Post-Handshake Certificate-Based Client Authentication
   (Section 4.6.2 of [TLS]) is performed after an Extended Key Update
   (EKU) is complete, the Handshake Context used for the transcript hash
   is updated.  It consists of transcript_hash_N+1 concatenated with the
   CertificateRequest message.  The Finished message is computed using a
   MAC key derived from the Base Key of the new epoch
   (client_application_traffic_secret_N+1).  This confirms that both
   peers are operating with the same updated traffic keys and completes
   an authenticated transition after the EKU.

   To prevent cryptographic state ambiguity, the following constraints
   apply:

   *  A TLS server MUST NOT initiate a post-handshake CertificateRequest
      while an EKU exchange is in progress.






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   *  If a CertificateRequest has been sent but the corresponding
      Finished message has not yet been received, the peers MUST NOT
      initiate an EKU exchange.

11.2.  Exported Authenticators

   This document updates Section 5.1 of [RFC9261] to specify that, after
   an Extended Key Update has completed, the Handshake Context and
   Finished MAC Key used for Exported Authenticators MUST be derived
   from the exporter secret associated with the current epoch.
   Implementations that support the epoch-aware Exported Authenticators
   interface MUST provide a means for applications to request the
   generation or validation of Exported Authenticators using the
   exporter secret for a specific epoch.

   The Handshake Context and Finished MAC Key used in both the
   CertificateVerify message (Section 5.2.2 of [RFC9261]) and the
   Finished message (Section 5.2.3 of [RFC9261]) are derived from the
   exporter secret associated with the current epoch.  If a MitM
   interferes with the EKU exchange and causes the peers to derive
   different traffic and exporter secrets, their Handshake Contexts and
   Finished MAC Keys will differ.  As a result, validation procedures
   specified in Section 5.2.4 of [RFC9261] will fail, thereby detecting
   the divergence of key state between peers.

   A new optional API SHOULD be defined to permit applications to
   request or verify Exported Authenticators for a specific exporter
   epoch.  As discussed in Section 7 of [RFC9261], this can, as an
   exception, be implemented at the application layer when the epoch-
   aware TLS exporter is available.  The APIs defined in [RFC9261]
   remain unchanged, so existing applications continue to operate
   without modification.  The epoch-aware API accepts an epoch
   identifier; when present, the (D)TLS implementation MUST derive the
   Handshake Context and Finished MAC Key from the exporter secret
   associated with that epoch.  When Exported Authenticators are
   generated using the epoch-aware Exported Authenticators interface,
   the epoch identifier used for their derivation can be conveyed in the
   certificate_request_context field, allowing the peer, particularly in
   DTLS where records may be reordered, to determine the correct
   exporter secret for validation.

12.  Security Considerations

   This section discusses additional security and operational aspects
   introduced by the Extended Key Update mechanism.  All security
   considerations of TLS 1.3 [TLS] and DTLS.13 [DTLS] continue to apply.





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12.1.  Scope of Key Compromise

   Extended Key Update assumes a transient compromise of the current
   application traffic keys, not a persistent attacker with ongoing
   access to key material.  The procedure itself does not rely on long-
   term private keys; those are assumed to remain secure, as they are
   typically stored in a secure element, such as a Trusted Execution
   Environment (TEE), Hardware Security Module (HSM), or Trusted
   Platform Module (TPM).  In contrast, application traffic keys are
   stored within the rich operating system, where short-term exposure
   due to memory disclosure or transient compromise may occur.  Post-
   compromise security can be re-established, provided the compromise is
   no longer active when an Extended Key Update is performed.

   Extended Key Update can restore confidentiality only if the attacker
   no longer has access to either peer.  If an adversary retains access
   to current application traffic keys and can act as a man-in-the-
   middle during the Extended Key Update, then the update cannot restore
   security unless Section 11 is used.

   If one of the mechanisms defined in Section 11 is not used, the
   attacker can impersonate each endpoint, substitute EKU messages, and
   maintain control of the communication.  When Post-handshake
   Certificate-Based Client Authentication or the modified Exported
   Authenticator mechanism is used, the authentication messages are
   bound to the keys established after the EKU.  Any modification or
   substitution of EKU messages therefore becomes detectable, preventing
   this attack.

   If a compromise occurs before the handshake completes, the ephemeral
   key exchange, client_handshake_traffic_secret,
   server_handshake_traffic_secret, and the initial client_/
   server_application_traffic_secret could be exposed.  In that case,
   only the initial handshake messages and the application data
   encrypted under the initial client_/server_application_traffic_secret
   can be decrypted until the Extended Key Update procedure completes.
   The Extended Key Update procedure derives fresh
   application_traffic_secrets from a new ephemeral key exchange,
   ensuring that all subsequent application data remains confidential.

12.2.  Post-Compromise Security

   Extended Key Update provides post-compromise security for long-lived
   TLS sessions.  To ensure post-compromise security guarantees:

   *  Each update MUST use freshly generated ephemeral key-exchange
      material.  Implementations MUST NOT reuse ephemeral key-exchange
      material across updates or across TLS sessions.



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12.3.  Denial-of-Service (DoS)

   The Extended Key Update mechanism increases computational and state-
   management overhead.  A malicious peer could attempt to exhaust CPU
   or memory resources by initiating excessive update requests.

   Implementations SHOULD apply the following mitigations:

   *  Limit the frequency of accepted Extended Key Update requests per
      session.

   *  A peer that has sent an Extended Key Update MUST NOT initiate
      another until the previous update completes.  If a peer violates
      this rule, the receiving peer MUST treat it as a protocol
      violation, send an "unexpected_message" alert, and terminate the
      connection.

12.4.  Operational Guidance

   Deployments SHOULD evaluate Extended Key Update performance under
   load and fault conditions, such as high-frequency or concurrent
   updates.  TLS policies SHOULD define explicit rate limits that
   balance post-compromise security benefits against potential DoS
   exposure.

13.  IANA Considerations

13.1.  TLS Flags

   IANA is requested to add the following entry to the "TLS Flags"
   extension registry [TLS-Ext-Registry]:

   *  Value: TBD1

   *  Flag Name: extended_key_update

   *  Messages: CH, EE

   *  Recommended: Y

   *  Reference: [This document]

13.2.  TLS HandshakeType

   IANA is requested to add the following entry to the "TLS
   HandshakeType" registry [TLS-Ext-Registry]:

   *  Value: TBD2



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   *  Description: extended_key_update

   *  DTLS-OK: Y

   *  Reference: [This document]

13.3.  ExtendedKeyUpdate Message Subtypes Registry

   IANA is requested to create a new registry "TLS ExtendedKeyUpdate
   Message Subtypes", within the existing "Transport Layer Security
   (TLS) Parameters" registry [TLS-Ext-Registry].  This new registry
   reserves types used for Extended Key Update entries.  The initial
   contents of this registry are as follows.

         +=======+=====================+=========+===============+
         | Value | Description         | DTLS-OK | Reference     |
         +=======+=====================+=========+===============+
         | 0     | key_update_request  | Y       | This document |
         +-------+---------------------+---------+---------------+
         | 1     | key_update_response | Y       | This document |
         +-------+---------------------+---------+---------------+
         | 2     | new_key_update      | Y       | This document |
         +-------+---------------------+---------+---------------+
         | 3-255 | Unassigned          |         |               |
         +-------+---------------------+---------+---------------+

                                  Table 1

   New assignments in the "TLS ExtendedKeyUpdate Types" registry will be
   administered by IANA through Specification Required procedure
   [RFC8126].  The role of the designated expert is described in
   Section 17 of [RFC8447].  The designated expert [RFC8126] ensures
   that the specification is publicly available.  It is sufficient to
   have an Internet-Draft (that is posted and never published as an RFC)
   or to cite a document from another standards body, industry
   consortium, or any other location.  An expert may provide more in-
   depth reviews, but their approval should not be taken as an
   endorsement of the ExtendedKeyUpdate Message Subtype.

13.4.  SSLKEYLOGFILE labels

   IANA is requested to add the following entries to the "TLS
   SSLKEYLOGFILE Labels" extension registry [TLS-Ext-Registry]:








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   +===========================+=============+===========+============+
   | Value                     | Description | Reference | Comment    |
   +===========================+=============+===========+============+
   | CLIENT_TRAFFIC_SECRET_N+1 | Secret      | This      | N          |
   |                           | protecting  | document  | represents |
   |                           | client      |           | iteration  |
   |                           | records     |           | of         |
   |                           | after       |           | Extended   |
   |                           | Extended    |           | Key Update |
   |                           | Key Update  |           |            |
   +---------------------------+-------------+-----------+------------+
   | SERVER_TRAFFIC_SECRET_N+1 | Secret      | This      | N          |
   |                           | protecting  | document  | represents |
   |                           | server      |           | iteration  |
   |                           | records     |           | of         |
   |                           | after       |           | Extended   |
   |                           | Extended    |           | Key Update |
   |                           | Key Update  |           |            |
   +---------------------------+-------------+-----------+------------+
   | EXPORTER_SECRET_N+1       | Exporter    | This      | N          |
   |                           | secret      | document  | represents |
   |                           | after       |           | iteration  |
   |                           | Extended    |           | of         |
   |                           | Key Update  |           | Extended   |
   |                           |             |           | Key Update |
   +---------------------------+-------------+-----------+------------+

                                 Table 2

14.  References

14.1.  Normative References

   [DTLS]     Rescorla, E., Tschofenig, H., and N. Modadugu, "The
              Datagram Transport Layer Security (DTLS) Protocol Version
              1.3", RFC 9147, DOI 10.17487/RFC9147, April 2022,
              <https://www.rfc-editor.org/rfc/rfc9147>.

   [RFC2119]  Bradner, S., "Key words for use in RFCs to Indicate
              Requirement Levels", BCP 14, RFC 2119,
              DOI 10.17487/RFC2119, March 1997,
              <https://www.rfc-editor.org/rfc/rfc2119>.

   [RFC3711]  Baugher, M., McGrew, D., Naslund, M., Carrara, E., and K.
              Norrman, "The Secure Real-time Transport Protocol (SRTP)",
              RFC 3711, DOI 10.17487/RFC3711, March 2004,
              <https://www.rfc-editor.org/rfc/rfc3711>.




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   [RFC5705]  Rescorla, E., "Keying Material Exporters for Transport
              Layer Security (TLS)", RFC 5705, DOI 10.17487/RFC5705,
              March 2010, <https://www.rfc-editor.org/rfc/rfc5705>.

   [RFC8174]  Leiba, B., "Ambiguity of Uppercase vs Lowercase in RFC
              2119 Key Words", BCP 14, RFC 8174, DOI 10.17487/RFC8174,
              May 2017, <https://www.rfc-editor.org/rfc/rfc8174>.

   [RFC9261]  Sullivan, N., "Exported Authenticators in TLS", RFC 9261,
              DOI 10.17487/RFC9261, July 2022,
              <https://www.rfc-editor.org/rfc/rfc9261>.

   [RFC9729]  Schinazi, D., Oliver, D., and J. Hoyland, "The Concealed
              HTTP Authentication Scheme", RFC 9729,
              DOI 10.17487/RFC9729, February 2025,
              <https://www.rfc-editor.org/rfc/rfc9729>.

   [TLS]      Rescorla, E., "The Transport Layer Security (TLS) Protocol
              Version 1.3", Work in Progress, Internet-Draft, draft-
              ietf-tls-rfc8446bis-14, 13 September 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-tls-
              rfc8446bis-14>.

   [TLS-FLAGS]
              Nir, Y., "A Flags Extension for TLS 1.3", Work in
              Progress, Internet-Draft, draft-ietf-tls-tlsflags-16, 14
              September 2025, <https://datatracker.ietf.org/doc/html/
              draft-ietf-tls-tlsflags-16>.

14.2.  Informative References

   [ANSSI]    ANSSI, "Recommendations for securing networks with IPsec,
              Technical Report", August 2015,
              <https://cyber.gouv.fr/sites/default/files/2012/09/
              NT_IPsec_EN.pdf>.

   [CCG16]    IEEE, "On Post-compromise Security", August 2016,
              <https://doi.org/10.1109/csf.2016.19>.

   [CONFIDENTIALITY]
              Barnes, R., Schneier, B., Jennings, C., Hardie, T.,
              Trammell, B., Huitema, C., and D. Borkmann,
              "Confidentiality in the Face of Pervasive Surveillance: A
              Threat Model and Problem Statement", RFC 7624,
              DOI 10.17487/RFC7624, August 2015,
              <https://www.rfc-editor.org/rfc/rfc7624>.





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   [RFC8126]  Cotton, M., Leiba, B., and T. Narten, "Guidelines for
              Writing an IANA Considerations Section in RFCs", BCP 26,
              RFC 8126, DOI 10.17487/RFC8126, June 2017,
              <https://www.rfc-editor.org/rfc/rfc8126>.

   [RFC8447]  Salowey, J. and S. Turner, "IANA Registry Updates for TLS
              and DTLS", RFC 8447, DOI 10.17487/RFC8447, August 2018,
              <https://www.rfc-editor.org/rfc/rfc8447>.

   [TLS-ECDHE-MLKEM]
              Kwiatkowski, K., Kampanakis, P., Westerbaan, B., and D.
              Stebila, "Post-quantum hybrid ECDHE-MLKEM Key Agreement
              for TLSv1.3", Work in Progress, Internet-Draft, draft-
              ietf-tls-ecdhe-mlkem-04, 8 February 2026,
              <https://datatracker.ietf.org/doc/html/draft-ietf-tls-
              ecdhe-mlkem-04>.

   [TLS-Ext-Registry]
              IANA, "Transport Layer Security (TLS) Extensions",
              November 2023, <https://www.iana.org/assignments/tls-
              extensiontype-values>.

   [TLS-HYBRID]
              Stebila, D., Fluhrer, S., and S. Gueron, "Hybrid key
              exchange in TLS 1.3", Work in Progress, Internet-Draft,
              draft-ietf-tls-hybrid-design-16, 7 September 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-tls-
              hybrid-design-16>.

   [TLS-KEYLOGFILE]
              Thomson, M., Rosomakho, Y., and H. Tschofenig, "The
              SSLKEYLOGFILE Format for TLS", Work in Progress, Internet-
              Draft, draft-ietf-tls-keylogfile-05, 9 June 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-tls-
              keylogfile-05>.

   [TLS-MLKEM]
              Connolly, D., "ML-KEM Post-Quantum Key Agreement for TLS
              1.3", Work in Progress, Internet-Draft, draft-ietf-tls-
              mlkem-07, 12 February 2026,
              <https://datatracker.ietf.org/doc/html/draft-ietf-tls-
              mlkem-07>.

   [TLS-RENEGOTIATION]
              Rescorla, E., Ray, M., Dispensa, S., and N. Oskov,
              "Transport Layer Security (TLS) Renegotiation Indication
              Extension", RFC 5746, DOI 10.17487/RFC5746, February 2010,
              <https://www.rfc-editor.org/rfc/rfc5746>.



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Appendix A.  Acknowledgments

   We would like to thank the members of the "TSVWG DTLS for SCTP
   Requirements Design Team" for their discussion.  The members, in no
   particular order, were:

   *  Marcelo Ricardo Leitner

   *  Zaheduzzaman Sarker

   *  Magnus Westerlund

   *  John Mattsson

   *  Claudio Porfiri

   *  Xin Long

   *  Michael Tüxen

   *  Hannes Tschofenig

   *  K Tirumaleswar Reddy

   *  Bertrand Rault

   Additionally, we would like to thank the chairs of the Transport and
   Services Working Group (tsvwg) Gorry Fairhurst and Marten Seemann as
   well as the responsible area director Martin Duke.

   Finally, we would like to thank Martin Thomson, Ilari Liusvaara,
   Benjamin Kaduk, Scott Fluhrer, Dennis Jackson, David Benjamin,
   Matthijs van Duin, Rifaat Shekh-Yusef, Joe Birr-Pixton, Eric
   Rescorla, and Thom Wiggers for their review comments.

Appendix B.  State Machines

   The sections below describe the state machines for the extended key
   update operation for TLS 1.3 and DTLS 1.3.

   The state machine diagrams in the Appendix are provided for
   illustrative purposes only to aid understanding of the protocol flow.
   They are not normative.  In case of any discrepancy between the
   Appendix diagrams and the protocol behavior specified in the main
   body of this document, the text in the draft takes precedence.

   For editorial reasons we abbreviate the protocol message types:




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   *  Req - ExtendedKeyUpdate(request)

   *  Resp - ExtendedKeyUpdate(response)

   *  NKU - ExtendedKeyUpdate(new_key_update)

   *  ACK - Acknowledgement message from Section 7 of [DTLS]

   *  APP - application data payloads

   In the (D)TLS 1.3 state machines discussed below, the terms SEND and
   RECEIVE keys refer to the send and receive key variables defined in
   Section 2.

   The TLS state machine diagrams use the following notation:

   *  Numbered labels (1), (2), … denote normal protocol transitions in
      the order they occur.

   *  Lettered labels (A), (B), … denote exceptional or conditional
      transitions.

   *  Self-loops indicate that the state does not change as a result of
      the event.

B.1.  TLS 1.3 State Machines

   This section describes the initiator and responder state machines.

B.1.1.  Initiator State Machine





















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   +----------------------+
   | START                |
   | send_key=current,    |
   | receive_key=current  |
   | updating=0           |
   +----------------------+
         |
   (1) send Req with local KeyShareEntry
       set updating=1
         v
   +----------------------+
   | WAIT_RESP            |
   | updating=1           |
   +----------------------+
   |  \
   |   \ (A) recv peer Req (crossed updates):
   |    \  compare peer.key_exchange to local.key_exchange
   |     \  - if peer < local: IGNORE peer Req (send nothing)
   |      \ - if peer == local: ABORT "unexpected_message"
   |       \ - if peer > local: ABANDON local update:
   |          * set updating=0
   |          * act as RESPONDER: send Resp with KeyShareEntry,
   |            derive new secrets, then proceed as in responder flow
   |
   |  (2) recv Resp with key_share:
   |      derive new secrets
   |      send NKU (encrypted under old keys)
   |      update SEND key (send_key := new)
   |      update RECEIVE key (receive_key := new)
   |      set updating=0
          v
   +----------------------+
   | FINISHED             |
   +----------------------+

   Notes:

   *  NKU message is sent under old keys.

   *  If a classic KeyUpdate arrives (EKU negotiated), ABORT
      "unexpected_message".

   *  Crossed-requests: ignore the request with LOWER lexicographic
      key_exchange; if equal, abort.

B.1.2.  Responder State Machine





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   +----------------------+
   | START                |
   | send_key=current,    |
   | receive_key=current  |
   | updating=0           |
   +----------------------+
         |
   (1) recv Req
       set updating=1
         v
   +----------------------+
   | RESPOND              |
   +----------------------+
   |  send Resp with KeyShareEntry
   |  (may defer sending if under load;
   |   must send once resources free)
   |  derive new secrets
   |  update SEND keys (send_key := new)
   |  --> WAIT_I_NKU
         v
   +----------------------+
   | WAIT_I_NKU           |
   | updating=1           |
   +----------------------+
         |
   (2) recv NKU (encrypted under old keys)
       update RECEIVE keys (receive_key := new)
       set updating=0
         v
   +----------------------+
   | FINISHED             |
   +----------------------+

   Notes:

   *  Responder may defer Resp under load (no status reply), then must
      send it once resources are free.

   *  If a classic KeyUpdate arrives (EKU negotiated), ABORT
      "unexpected_message".

B.2.  DTLS 1.3 State Machines

   This section describes the initiator and responder state machines.







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B.2.1.  Terms and Abbreviations

   The following variables and abbreviations are used in the state
   machine diagrams.

   *  rx - current, accepted receive epoch.

   *  tx - current transmit epoch used for tagging outgoing messages.

   *  E - initial epoch value.

   *  updating - true while a key-update handshake is in progress.

   *  old_rx - the previous receive epoch remembered during retention.

   *  retain_old - when true, receiver accepts tags old_rx and rx.

   *  tag=... - the TX-epoch value written on an outgoing message.

   *  e==... - the epoch tag carried on an incoming message (what the
      peer sent).

   *  FINISHED / START / WAIT_RESP / WAIT_I_NKU / WAIT_R_NKU / ACTIVATE
      RETENTION / RESPOND / WAIT_ACK - diagram states; FINISHED denotes
      the steady state after success.

   Crossed requests.  If both peers independently initiate the extended
   key update and the key_update_request messages cross in flight,
   compare the KeyShareEntry.key_exchange values.  The request with the
   lower lexicographic value must be ignored.  If the values are equal,
   the endpoint must abort with an "unexpected_message" alert.  If the
   peer's value is higher than the local one, the endpoint abandons its
   in-flight update and processes the peer's request as responder.

B.2.2.  State Machine (Initiator)

   The initiator starts in the START state with matching epochs (rx :=
   E; tx := E).  It sends a Req and enters WAIT_RESP (updating := 1).
   While waiting, APP data may be sent at any time (tagged with the
   current tx) and received according to the APP acceptance rule below.

   Once the responder returns Resp with a tag matching the current rx,
   the initiator derives new key material.  It then sends NKU still
   tagged with the old tx.  The initiator activates retention mode: the
   old epoch is remembered, the receive epoch is incremented, and
   application data is accepted under both epochs for a transition
   period.  Initiator moves to WAIT_ACK.




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   If a peer key_update_request arrives while in WAIT_RESP (crossed
   updates), apply the crossed-request rule above.  If the peer's
   key_exchange is higher, abandon the local update (updating := 0) and
   continue as responder: send key_update_response, derive new secrets,
   then proceed with the responder flow.  If lower, ignore the peer's
   request; if equal, abort with "unexpected_message".

   Upon receiving the responder's ACK matching the updated epoch, the
   initiator completes the transition by synchronizing transmit and
   receive epochs (tx := rx), disabling retention, and clearing the
   update flag.  The state machine returns to FINISHED, ready for
   subsequent updates.

   Throughout the process:

   *  Duplicate messages are tolerated (for retransmission handling).

   *  Temporary epoch mismatches are permitted while an update is in
      progress.

   *  Application data flows continuously, subject to epoch acceptance
      rules.

   +---------------------------------+
   | START                           |
   | rx := E; tx := E, updating := 0 |
   +---------------------------------+
         |
   (1) send Req [tag=tx]
       set updating := 1
         v
   +------------------------------+
   | WAIT_RESP                    |
   +------------------------------+
    ^    ^    ^           ^
    |    |    |           |
    |    |    |           |
    |    |    |           +-- (A) recv peer Req (crossed):
    |    |    |                if peer < local  -> IGNORE (Self-Loop)
    |    |    |                if peer == local -> ABORT  (Error)
    |    |    |                if peer > local  -> ABANDON update;
    |    |    |                                    updating := 0;
    |         |    |                                    act as RESPONDER
    |    |    |
    |    |    +-- APP send (anytime)           (Self-Loop)
    |    |
    |    +-- APP recv: retain_old && e == old_rx   (Self-Loop)
    |



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    +-- APP recv: e == rx                           (Self-Loop)

             | recv Resp [e == rx]
             | derive secrets
             v
   +---------------------+
   | ACTIVATE RETENTION  |
   | old_rx=rx;          |
   | retain_old=1; rx++  |
   +---------------------+
             |
   (2) send NKU [tag=old tx]
             v
   +------------------------------+
   | WAIT_ACK (updating = 1)      |
   +------------------------------+
         |
         |  APP send/recv allowed
         |
   (3) recv ACK [e==rx]
       tx=rx; retain_old=0; updating := 0
         v
   +------------------------------------------+
   | FINISHED                                 |
   +------------------------------------------+

B.2.3.  State Machine (Responder)

   The responder starts in the START state with synchronized transmit
   and receive epochs (rx := E; tx := E) and no update in progress.
   Application data can be transmitted at any time using the sender's
   current transmit epoch.  A receiver must accept application data if
   the epoch tag on the DTLS record equals the receiver's current
   receive epoch.  If the receiver has retention active (retain_old ==
   true), the receiver must also accept DTLS records whose epoch tag
   equals the remembered previous epoch.

   Upon receiving an ExtendedKeyUpdate(key_update_request) (Req), the
   responder transitions to RESPOND.  The responder may defer sending
   ExtendedKeyUpdate(key_update_response) under load; in that case it
   must acknowledge the request with an ACK and retransmit the response
   until it is acknowledged by the initiator, as specified in DTLS
   considerations.  When sent, ExtendedKeyUpdate(key_update_response) is
   tagged with the current tx.

   After an ExtendedKeyUpdate(new_key_update) (NKU) is received with the
   correct epoch, the responder:




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   1.  activates retention (old_rx := rx; retain_old := 1),

   2.  increments both epochs (rx++, tx++),

   3.  sends ACK tagged with the new tx (which now equals the new rx),

   4.  clears updating and enters FINISHED.

   Retention at the responder ends automatically on the first APP
   received under the new rx (then retain_old := 0).  APP traffic is
   otherwise permitted at any time; reordering is tolerated by the
   acceptance rule.

   APP acceptance rule (receiver): accept if e == rx or (retain_old && e
   == old_rx).  If retain_old is set and an APP with the new rx arrives,
   clear retain_old.



































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   +---------------------------------+
   | START                           |
   | rx := E; tx := E, updating := 0 |
   +---------------------------------+
         |
   (1) recv Req [e==rx]
       set updating := 1
         v
   +----------------------+
   | RESPOND              |
   +----------------------+
         |
         | send Resp [tag=tx]
         | (may defer; if deferred, ACK Req and later retransmit Resp
         |  until acknowledged by the initiator)
         v
   +---------------+
   | WAIT_I_NKU    |
   | (updating=1)  |
   +-------+-------+
         |
   (2) recv NKU [e==rx]      (assert accepted)
         |
   +---------------------+
   | ACTIVATE RETENTION  |
   | old_rx=rx;          |
   | retain_old=1;       |
   | rx=rx+1; tx=tx+1    |
   +----------+----------+
         |
   (3) send ACK [tag=tx]
       set updating := 0; assert tx==rx
         v
   +------------------------------------------+
   | FINISHED                                 |
   | (retain_old=0 after first APP at new rx) |
   +------------------------------------------+

Appendix C.  Overview of Security Goals

   A complete security analysis of the EKU is outside the scope of this
   document.  This appendix provides an informal description of the
   primary security goals that EKU is designed to achieve.








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C.1.  Post-Compromise Security (PCS)

   Extended Key Update supports post-compromise security under the
   assumptions described in Section 12.1.  If an attacker temporarily
   compromises an endpoint and obtains the traffic keys in use before an
   Extended Key Update takes place, but the compromise does not persist
   after the EKU completes, the attacker cannot derive the new keying
   material established by EKU.  This property follows from the use of
   fresh ephemeral key exchange material during each Extended Key
   Update, which produces new traffic keys that are independent of the
   previous ones.  This property provides only best-effort post-
   compromise security, as it assumes the attacker is not acting as a
   MiTM during the Extended Key Update.

   As a result, confidentiality of application data encrypted after the
   Extended Key Update is preserved even if the earlier traffic keys
   were exposed.

C.2.  Key Freshness and Cryptographic Independence

   Each Extended Key Update derives new traffic keys from ephemeral key
   exchange material.  This ensures strong separation between successive
   traffic keys:

   *  The new traffic keys established by an Extended Key Update are
      independent of the previous traffic keys.

   *  Compromise of one of traffic keys does not allow recovery of any
      earlier or later traffic keys.

   *  Application data protected under one of the traffic keys cannot be
      decrypted using keys from another.

C.3.  Elimination of Standard KeyUpdate

   Once Extended Key Update has been negotiated for a session, peers
   rely exclusively on EKU rather than the standard TLS 1.3 KeyUpdate
   mechanism.  Relying solely on Extended Key Update helps maintain PCS
   properties throughout the lifetime of the TLS session.

C.4.  Detecting Divergent Key State

   As described in Section 11, both Post-handshake Certificate-Based
   Client Authentication and Exported Authenticators can be used after
   an Extended Key Update to confirm that both endpoints derived the
   same traffic keys.  Because the authentication messages produced by
   these mechanisms depend on values derived from the updated traffic
   keys, any divergence in those traffic keys causes validation to fail,



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   revealing interference by an active attacker.

Authors' Addresses

   Hannes Tschofenig
   Siemens
   Email: hannes.tschofenig@gmx.net


   Michael Tüxen
   Münster Univ. of Applied Sciences
   Email: tuexen@fh-muenster.de


   Tirumaleswar Reddy
   Nokia
   Email: kondtir@gmail.com


   Steffen Fries
   Siemens
   Email: steffen.fries@siemens.com


   Yaroslav Rosomakho
   Zscaler
   Email: yrosomakho@zscaler.com
























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name: draft-ietf-anima-constrained-voucher
rev: 30
title: Constrained Bootstrapping Remote Secure Key Infrastructure (cBRSKI)
time: 2026-02-27T12:49:46Z
group_uri: /api/v1/group/group/1956/
expires: 2026-08-31T12:49:46Z
shepherd: /api/v1/person/email/tte@cs.fau.de/
states: ["/api/v1/doc/state/1/", "/api/v1/doc/state/150/", "/api/v1/doc/state/41/"]
categories: ["Agent identity/auth", "Data formats/interop"]
tags: []
fetched_at: 2026-02-27T22:56:09.014969+00:00
abstract:
This document defines the Constrained Bootstrapping Remote Secure Key
   Infrastructure (cBRSKI) protocol, which provides a solution for
   secure zero-touch onboarding of resource-constrained (IoT) devices
   into the network of a domain owner.  This protocol is designed for
   constrained networks, which may have limited data throughput or may
   experience frequent packet loss. cBRSKI is a variant of the BRSKI
   protocol, which uses an artifact signed by the device manufacturer
   called the "voucher" which enables a new device and the owner's
   network to mutually authenticate.  While the BRSKI voucher data is
   encoded in JSON, cBRSKI uses a compact CBOR-encoded voucher.  The
   BRSKI voucher data definition is extended with new data types that
   allow for smaller voucher sizes.  The Enrollment over Secure
   Transport (EST) protocol, used in BRSKI, is replaced with EST-over-
   CoAPS; and HTTPS used in BRSKI is replaced with DTLS-secured CoAP
   (CoAPS).  This document Updates RFC 8995 and RFC 9148.

full_text:
anima Working Group                                        M. Richardson
Internet-Draft                                  Sandelman Software Works
Updates: 8995, 9148 (if approved)                        P. van der Stok
Intended status: Standards Track                  vanderstok consultancy
Expires: 31 August 2026                                    P. Kampanakis
                                                           Cisco Systems
                                                                 E. Dijk
                                                       IoTconsultancy.nl
                                                        27 February 2026


  Constrained Bootstrapping Remote Secure Key Infrastructure (cBRSKI)
                draft-ietf-anima-constrained-voucher-30

Abstract

   This document defines the Constrained Bootstrapping Remote Secure Key
   Infrastructure (cBRSKI) protocol, which provides a solution for
   secure zero-touch onboarding of resource-constrained (IoT) devices
   into the network of a domain owner.  This protocol is designed for
   constrained networks, which may have limited data throughput or may
   experience frequent packet loss. cBRSKI is a variant of the BRSKI
   protocol, which uses an artifact signed by the device manufacturer
   called the "voucher" which enables a new device and the owner's
   network to mutually authenticate.  While the BRSKI voucher data is
   encoded in JSON, cBRSKI uses a compact CBOR-encoded voucher.  The
   BRSKI voucher data definition is extended with new data types that
   allow for smaller voucher sizes.  The Enrollment over Secure
   Transport (EST) protocol, used in BRSKI, is replaced with EST-over-
   CoAPS; and HTTPS used in BRSKI is replaced with DTLS-secured CoAP
   (CoAPS).  This document Updates RFC 8995 and RFC 9148.

About This Document

   This note is to be removed before publishing as an RFC.

   Status information for this document may be found at
   https://datatracker.ietf.org/doc/draft-ietf-anima-constrained-
   voucher/.

   Discussion of this document takes place on the anima Working Group
   mailing list (mailto:anima@ietf.org), which is archived at
   https://mailarchive.ietf.org/arch/browse/anima/.  Subscribe at
   https://www.ietf.org/mailman/listinfo/anima/.

   Source for this draft and an issue tracker can be found at
   https://github.com/anima-wg/constrained-voucher.




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Status of This Memo

   This Internet-Draft is submitted in full conformance with the
   provisions of BCP 78 and BCP 79.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF).  Note that other groups may also distribute
   working documents as Internet-Drafts.  The list of current Internet-
   Drafts is at https://datatracker.ietf.org/drafts/current/.

   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet-Drafts as reference
   material or to cite them other than as "work in progress."

   This Internet-Draft will expire on 31 August 2026.

Copyright Notice

   Copyright (c) 2026 IETF Trust and the persons identified as the
   document authors.  All rights reserved.

   This document is subject to BCP 78 and the IETF Trust's Legal
   Provisions Relating to IETF Documents (https://trustee.ietf.org/
   license-info) in effect on the date of publication of this document.
   Please review these documents carefully, as they describe your rights
   and restrictions with respect to this document.  Code Components
   extracted from this document must include Revised BSD License text as
   described in Section 4.e of the Trust Legal Provisions and are
   provided without warranty as described in the Revised BSD License.

Table of Contents

   1.  Introduction  . . . . . . . . . . . . . . . . . . . . . . . .   5
   2.  Terminology . . . . . . . . . . . . . . . . . . . . . . . . .   6
   3.  Requirements Language . . . . . . . . . . . . . . . . . . . .   7
   4.  Overview of Protocol  . . . . . . . . . . . . . . . . . . . .   7
   5.  Updates to RFC 8995 and RFC 9148  . . . . . . . . . . . . . .   9
   6.  BRSKI-EST Protocol  . . . . . . . . . . . . . . . . . . . . .  10
     6.1.  DTLS Connection . . . . . . . . . . . . . . . . . . . . .  10
       6.1.1.  DTLS Version  . . . . . . . . . . . . . . . . . . . .  10
       6.1.2.  TLS Client Certificates: IDevID authentication  . . .  10
       6.1.3.  DTLS Handshake Fragmentation Considerations . . . . .  11
       6.1.4.  Registrar and the Server Name Indicator (SNI) . . . .  11
       6.1.5.  Registrar Server Certificate Requirements . . . . . .  12
     6.2.  cBRSKI Join Proxy . . . . . . . . . . . . . . . . . . . .  12
     6.3.  Request URIs, Resource Discovery and Content-Formats  . .  12
       6.3.1.  Status Telemetry Returns  . . . . . . . . . . . . . .  14



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       6.3.2.  CoAP Resources Table  . . . . . . . . . . . . . . . .  15
       6.3.3.  CoAP Uri-Path Abbreviation  . . . . . . . . . . . . .  15
     6.4.  CoAP Responses  . . . . . . . . . . . . . . . . . . . . .  15
     6.5.  Extensions to EST-coaps . . . . . . . . . . . . . . . . .  16
       6.5.1.  Pledge Enrollment Procedure . . . . . . . . . . . . .  16
       6.5.2.  Renewal of CA certificates  . . . . . . . . . . . . .  17
       6.5.3.  Change of Domain Trust Anchor(s)  . . . . . . . . . .  17
       6.5.4.  Re-enrollment Procedure . . . . . . . . . . . . . . .  18
       6.5.5.  Multipart Content-Format for CA certificates (/crts)
               Resource  . . . . . . . . . . . . . . . . . . . . . .  19
     6.6.  Registrar Extensions  . . . . . . . . . . . . . . . . . .  20
   7.  BRSKI-MASA Protocol . . . . . . . . . . . . . . . . . . . . .  21
     7.1.  Protocol and Formats  . . . . . . . . . . . . . . . . . .  21
     7.2.  Registrar Voucher Request . . . . . . . . . . . . . . . .  22
     7.3.  MASA and the Server Name Indicator (SNI)  . . . . . . . .  22
     7.4.  Registrar Client Certificate Requirements . . . . . . . .  23
   8.  Pinning in Voucher Artifacts  . . . . . . . . . . . . . . . .  23
     8.1.  Registrar Identity Selection and Encoding . . . . . . . .  23
     8.2.  MASA Pinning Policy . . . . . . . . . . . . . . . . . . .  24
     8.3.  Pinning of Raw Public Keys (RPK)  . . . . . . . . . . . .  25
   9.  Artifacts . . . . . . . . . . . . . . . . . . . . . . . . . .  26
     9.1.  Example Artifacts . . . . . . . . . . . . . . . . . . . .  27
       9.1.1.  Example Pledge Voucher Request (PVR) Artifact . . . .  27
       9.1.2.  Example Registrar Voucher Request (RVR) Artifact  . .  27
       9.1.3.  Example Voucher Artifacts . . . . . . . . . . . . . .  28
     9.2.  Signing Voucher and Voucher Request Artifacts with
           COSE  . . . . . . . . . . . . . . . . . . . . . . . . . .  29
       9.2.1.  Signing of Registrar Voucher Request (RVR)  . . . . .  30
       9.2.2.  Signing of Pledge Voucher Request (PVR) . . . . . . .  31
       9.2.3.  Signing of Voucher by MASA  . . . . . . . . . . . . .  32
       9.2.4.  Optional Validation of Voucher by Registrar . . . . .  34
       9.2.5.  Additional Information in the COSE Header . . . . . .  34
   10. Extensions to Discovery . . . . . . . . . . . . . . . . . . .  35
     10.1.  Discovery Operations by a Pledge . . . . . . . . . . . .  36
       10.1.1.  Examples . . . . . . . . . . . . . . . . . . . . . .  37
     10.2.  Discovery Operations by a Join Proxy . . . . . . . . . .  39
   11. Deployment-specific Discovery Considerations  . . . . . . . .  39
     11.1.  6TiSCH Deployments . . . . . . . . . . . . . . . . . . .  39
     11.2.  IP networks using GRASP  . . . . . . . . . . . . . . . .  39
     11.3.  IP networks using mDNS . . . . . . . . . . . . . . . . .  40
     11.4.  Thread Networks using Mesh Link Establishment (MLE)  . .  40
   12. Design and Implementation Considerations  . . . . . . . . . .  41
     12.1.  Voucher Format and Encoding  . . . . . . . . . . . . . .  41
     12.2.  CoAP Usage . . . . . . . . . . . . . . . . . . . . . . .  41
     12.3.  Use of cBRSKI with HTTPS . . . . . . . . . . . . . . . .  41
   13. Raw Public Key Variant  . . . . . . . . . . . . . . . . . . .  42
     13.1.  Introduction and Scope . . . . . . . . . . . . . . . . .  42
     13.2.  DTLS Connection and Registrar Trust Anchor . . . . . . .  42



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     13.3.  The Pledge Voucher Request . . . . . . . . . . . . . . .  43
     13.4.  The Voucher Response . . . . . . . . . . . . . . . . . .  43
     13.5.  The Enrollment Phase . . . . . . . . . . . . . . . . . .  44
   14. Security Considerations . . . . . . . . . . . . . . . . . . .  44
     14.1.  Duplicate Serial Numbers . . . . . . . . . . . . . . . .  44
     14.2.  IDevID Security in the Pledge  . . . . . . . . . . . . .  45
     14.3.  Security of the BRSKI-MASA Protocol  . . . . . . . . . .  46
     14.4.  Registrar Certificate May Be Self-signed . . . . . . . .  47
     14.5.  Use of RPK Alternatives to 'proximity-registrar-cert'  .  47
   15. IANA Considerations . . . . . . . . . . . . . . . . . . . . .  48
     15.1.  Resource Type Link Target Attribute Values Registry  . .  48
     15.2.  Media Types Registry . . . . . . . . . . . . . . . . . .  48
       15.2.1.  application/voucher+cose . . . . . . . . . . . . . .  48
       15.2.2.  Interoperability Considerations for application/
               voucher+cose  . . . . . . . . . . . . . . . . . . . .  49
     15.3.  CoAP Content-Formats Registry  . . . . . . . . . . . . .  50
     15.4.  Update to BRSKI Well-Known URIs Registry . . . . . . . .  50
     15.5.  Structured Syntax Suffixes Registry  . . . . . . . . . .  51
   16. References  . . . . . . . . . . . . . . . . . . . . . . . . .  52
     16.1.  Normative References . . . . . . . . . . . . . . . . . .  52
     16.2.  Informative References . . . . . . . . . . . . . . . . .  56
   Appendix A.  Software and Library Support for cBRSKI  . . . . . .  58
     A.1.  Open Source cBRSKI Implementations  . . . . . . . . . . .  58
     A.2.  Security Library Support  . . . . . . . . . . . . . . . .  59
       A.2.1.  OpensSSL Example Code . . . . . . . . . . . . . . . .  60
       A.2.2.  mbedTLS Example Code  . . . . . . . . . . . . . . . .  61
     A.3.  Generating Certificates with OpenSSL  . . . . . . . . . .  61
   Appendix B.  cBRSKI Message Examples  . . . . . . . . . . . . . .  65
     B.1.  enrollstatus  . . . . . . . . . . . . . . . . . . . . . .  65
     B.2.  voucher_status  . . . . . . . . . . . . . . . . . . . . .  67
   Appendix C.  COSE-signed Voucher (Request) Examples . . . . . . .  68
     C.1.  Pledge, Registrar and MASA Keys . . . . . . . . . . . . .  68
       C.1.1.  Pledge IDevID Private Key . . . . . . . . . . . . . .  68
       C.1.2.  Registrar Private Key . . . . . . . . . . . . . . . .  68
       C.1.3.  MASA Private Key  . . . . . . . . . . . . . . . . . .  69
     C.2.  Pledge, Registrar, Domain CA and MASA Certificates  . . .  69
       C.2.1.  Pledge IDevID Certificate . . . . . . . . . . . . . .  69
       C.2.2.  Registrar Certificate . . . . . . . . . . . . . . . .  71
       C.2.3.  Domain CA Certificate . . . . . . . . . . . . . . . .  73
       C.2.4.  MASA Certificate  . . . . . . . . . . . . . . . . . .  75
     C.3.  COSE-signed Pledge Voucher Request (PVR)  . . . . . . . .  77
     C.4.  COSE-signed Registrar Voucher Request (RVR) . . . . . . .  78
     C.5.  COSE-signed Voucher from MASA . . . . . . . . . . . . . .  81
   Appendix D.  Pledge Device Class Profiles . . . . . . . . . . . .  83
     D.1.  Minimal Pledge  . . . . . . . . . . . . . . . . . . . . .  83
     D.2.  Typical Pledge  . . . . . . . . . . . . . . . . . . . . .  83
     D.3.  Full-featured Pledge  . . . . . . . . . . . . . . . . . .  84
     D.4.  Comparison Chart of Pledge Classes  . . . . . . . . . . .  84



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   Appendix E.  Pledge Discovery of Onboarding and Enrollment
           Options . . . . . . . . . . . . . . . . . . . . . . . . .  86
     E.1.  Pledge Discovery Query for All cBRSKI Resources . . . . .  86
     E.2.  Pledge Discovery Query for the cBRSKI Base Resource . . .  88
     E.3.  Usage of ct Attribute . . . . . . . . . . . . . . . . . .  88
     E.4.  EST-coaps Resource Discovery  . . . . . . . . . . . . . .  89
   Acknowledgements  . . . . . . . . . . . . . . . . . . . . . . . .  90
   Changelog . . . . . . . . . . . . . . . . . . . . . . . . . . . .  91
   Authors' Addresses  . . . . . . . . . . . . . . . . . . . . . . .  93

1.  Introduction

   Secure enrollment of new nodes into constrained networks with
   constrained nodes presents unique challenges.  As explained in
   [RFC7228], such networks may have limited data throughput or may
   experience frequent packet loss.  In addition, its nodes may be
   constrained by energy availability, memory space, and code size.

   The Bootstrapping Remote Secure Key Infrastructure (BRSKI) protocol
   described in [RFC8995] provides a solution for secure zero-touch
   (automated) onboarding of new (unconfigured) devices.  These new
   devices are called "Pledges", equipped with a factory-installed
   Initial Device Identifier (IDevID) (see [ieee802-1AR]).  Using the
   IDevID, a Pledge is securely enrolled into a network.

   The BRSKI solution described in [RFC8995] was designed to be modular,
   and this document describes a version scaled to the constraints of
   IoT deployments.  This document uses the constrained voucher and
   voucher request artifacts defined in [RFC8366bis] for a constrained
   version of the BRSKI protocol: cBRSKI.  The cBRSKI protocol uses the
   CoAP-based version of EST (EST-coaps from [RFC9148]) rather than the
   EST over HTTPS [RFC7030].  cBRSKI is itself scalable to multiple
   resource levels through the definition of optional functions.
   Appendix D illustrates this.

   In BRSKI, the [RFC8366bis] voucher data is by default serialized to
   JSON with a signature in CMS [RFC5652]. cBRSKI uses the CBOR
   [RFC8949] voucher data serialization defined by [RFC8366bis], and
   applies a new COSE [RFC9052] signature format as defined in
   Section 9.

   This COSE-signed CBOR-encoded voucher is transported using both
   secured CoAP [RFC7252] and HTTPS.  The CoAP connection (between
   Pledge and Registrar) is to be protected by DTLS (CoAPS).  The HTTP
   connection (between Registrar and MASA) is to be protected using TLS
   (HTTPS).





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   Section 4 to Section 10 define the default cBRSKI protocol, by means
   of additions to and modifications of regular BRSKI.  Section 11
   considers some variations of the protocol, specific to particular
   deployments or IoT networking technologies.  Next in Section 12, some
   considerations for the design and implementation of cBRSKI components
   are provided.

   Section 13 introduces a variant of cBRSKI for the most-constrained
   Pledges, using Raw Public Keys (RPK).  This variant achieves smaller
   sizes of data objects and avoids doing certain costly PKIX
   verification operations on the Pledge.

   Appendix E provides more details on how a Pledge may discover the
   various onboarding/enrollment options that a Registrar provides.
   Implementing these methods is optional for a Pledge.

2.  Terminology

   The following terms are defined in [RFC8366bis], and are used
   identically as in that document: Artifact, Attribute, Domain, Join
   Registrar and Coordinator (JRC), Malicious Registrar, Manufacturer
   Authorized Signing Authority (MASA), Pledge, Registrar, Onboarding,
   Owner, Voucher Data, Voucher Request and Voucher.

   The protocol described in this document is referred to as cBRSKI, the
   constrained version of BRSKI [RFC8995].

   The following terms from [RFC8995] are used identically as in that
   document: Domain CA, enrollment, IDevID, Join Proxy, LDevID,
   manufacturer, nonced, nonceless, PKIX.

   The following terms from [RFC7030] are used identically as in that
   document: Explicit Trust Anchor (TA), Explicit TA database, Third-
   party TA.

   The following terms from [RFC7252] are used identically as in that
   document: Confirmable (CON), Acknowledgement (ACK), Endpoint, ETag,
   Client, Server, Piggybacked Response, resource, Resource Discovery,
   Content-Format.

   The term Pledge Voucher Request, or acronym PVR, is introduced to
   refer to the voucher request between the Pledge and the Registrar.

   The term Registrar Voucher Request, or acronym RVR, is introduced to
   refer to the voucher request between the Registrar and the MASA.

   The terms "PKIX Certificate" and "certificate" both refer to the
   X.509v3 profile described in [RFC5280].



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   The term "base resource" is defined as a CoAP resource that can be
   used as a base to append an additional path segment to, where this
   segment is a short resource name ('short-name') as defined in
   Section 6.3 and Table 1.

   In code examples, the string "<CODE BEGINS>" denotes the start of a
   code example and "<CODE ENDS>" the end of the code example. "lf
   added" means that extra linefeed characters were added to an example
   to make lines fit in this document.

   The ellipsis ("...") in a CBOR diagnostic notation byte string
   denotes a further sequence of bytes that is not shown for brevity.
   This notation is defined in [I-D.ietf-cbor-edn-literals].

3.  Requirements Language

   The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT",
   "SHOULD", "SHOULD NOT", "RECOMMENDED", "NOT RECOMMENDED", "MAY", and
   "OPTIONAL" in this document are to be interpreted as described in
   BCP 14 [RFC2119] [RFC8174] when, and only when, they appear in all
   capitals, as shown here.

4.  Overview of Protocol

   [RFC8366bis] defines a voucher that can assert proximity,
   authenticates the Registrar, and can offer varying levels of anti-
   replay protection.  The proximity proof provided by a voucher is an
   assertion that the Pledge and the Registrar are believed to be close
   together, from a network topology point of view.  Similar to BRSKI
   [RFC8995], proximity is proven by making a DTLS connection between a
   Pledge and a Registrar.  The Pledge initiates this connection using a
   link-local source address.

   The secure DTLS connection is then used by the Pledge to send a
   Pledge Voucher Request (PVR).  The Registrar then includes the PVR
   into its own Registrar Voucher Request (RVR), which is sent to an
   agent (MASA) of the Pledge's manufacturer.  The MASA verifies the PVR
   and RVR and issues a signed voucher.  The voucher provides an
   authorization statement from the manufacturer indicating that the
   Registrar is the intended owner of the Pledge.  The voucher refers to
   the Registrar through pinning of the Registrar's identity.










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   After verification of the voucher, the Pledge enrolls into the
   Registrar's domain by obtaining a certificate using the EST-coaps
   [RFC9148] protocol, suitable for constrained devices.  Once the
   Pledge has obtained its domain identity (LDevID) in this manner, it
   can use this identity to obtain network access credentials, which are
   used to join the local IP network.  The method to obtain such
   credentials depends on the particular network technology used and is
   outside the scope of this document.

   The two main parts of the BRSKI protocol are named separately in this
   document: BRSKI-EST (Section 6) for the protocol between Pledge and
   Registrar, and BRSKI-MASA (Section 7) for the protocol between the
   Registrar and the MASA.

   Time-based vouchers are supported, but given that constrained devices
   are unlikely to have accurate time, their use will be uncommon.  Most
   Pledges using constrained vouchers will be online during enrollment
   and will use live nonces to provide anti-replay protection rather
   than expiry times.

   [RFC8366bis] defines the CBOR voucher data encoding for the
   constrained voucher and the constrained voucher request, which are
   used by cBRSKI.

   The constrained voucher request MUST be signed by the Pledge.  COSE
   [RFC9052] is used for signing as defined in Section 9.2.  It signs
   using the private key of its IDevID.  The constrained voucher MUST be
   signed by the MASA.  Also in this case, COSE is used for signing.

   For the constrained voucher request (PVR) the default method for the
   Pledge to identify the Registrar is using the Registrar's full PKIX
   certificate.  But when operating PKIX-less as described in
   Section 13, the Registrar's Raw Public Key (RPK) is used for this.

   For the constrained voucher the default method to indicate ("pin") a
   trusted domain identity is the domain's PKIX CA certificate, but when
   operating PKIX-less instead the RPK of the Registrar is pinned.

   For certificates, cBRSKI currently uses the X.509 format, like BRSKI.
   The protocol and data formats are defined such that future extension
   to other certificate formats is enabled.  For example, CBOR-encoded
   and COSE-signed C509 certificates ([I-D.ietf-cose-cbor-encoded-cert])
   may provide data size savings as well as code sharing benefits with
   CBOR/COSE libraries, when applied to cBRSKI.

   The BRSKI architecture mandates that the MASA be aware of the
   capabilities of the Pledge.  This is not a drawback as a Pledge is
   constructed by a manufacturer which also arranges for the MASA to be



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   aware of the inventory of devices.  The MASA therefore knows if the
   Pledge supports PKIX operations, or if it is limited to RPK
   operations only.  Based upon this, the MASA can select which
   attributes to use in the voucher for certain operations, like the
   pinning of the Registrar or domain identity.

5.  Updates to RFC 8995 and RFC 9148

   This section details the ways in which this document updates other
   RFCs.

   This document Updates [RFC8995] because it:

   *  clarifies how pinning in vouchers is done (Section 8),

   *  clarifies the use of TLS Server Name Indicator (SNI)
      (Section 6.1.4, Section 7.3),

   *  clarifies when new trust anchors should be retrieved by a Pledge
      (Section 6.5.1),

   *  clarifies what kinds of Extended Key Usage attributes are
      appropriate for each certificate (Section 6.1.5, Section 7.4),

   *  extends BRSKI with the use of CoAP,

   *  makes some BRSKI messages optional to send if the results can be
      inferred from other validations (Section 6.5),

   *  extends the BRSKI-EST/BRSKI-MASA protocols (Section 6, Section 7,
      Section 9.2) to carry the new application/voucher+cose format.

   This document Updates [RFC9148] because it:

   *  defines stricter DTLS requirements (Section 6.1)),

   *  details how an EST-coaps client handles certificate renewal and
      re-enrollment (Section 6.5),

   *  details how an EST-coaps server processes a "CA certificates"
      request for content-format 287 (application/pkix-cert)
      (Section 6.6).

   *  adds enrollment status telemetry to the certificate renewal
      procedure (Section 6.5.4),

   *  adds support for the media type application/multipart-core for the
      CA certificates (/crts) resource (Section 6.5.5),



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   *  defines a resource type ('rt') attribute value "ace.est" for the
      EST-coaps base resource (Section 15.1).

6.  BRSKI-EST Protocol

   This section describes the extensions to both BRSKI [RFC8995] and
   EST-coaps [RFC9148] operations between Pledge and Registrar.

6.1.  DTLS Connection

   A DTLS connection is established between the Pledge and the
   Registrar, similar to the TLS connection described in Section 5.1 of
   [RFC8995].  This may occur via a Join Proxy as described in
   Section 6.2.  Regardless of the Join Proxy presence or particular
   mechanism used, the DTLS connection should operate identically.  The
   cBRSKI and EST-coaps requests and responses for onboarding are
   carried over this DTLS connection.

6.1.1.  DTLS Version

   DTLS version 1.3 [RFC9147] SHOULD be used in any implementation of
   this specification.  An exception case where DTLS 1.2 [RFC6347] MAY
   be used is in a Pledge that uses a software platform where a DTLS 1.3
   client is not available (yet).  This may occur for example if a
   legacy device gets software-upgraded to support cBRSKI.  For this
   reason, a Registrar MUST by default support both DTLS 1.3 and DTLS
   1.2 client connections.  However, for security reasons the Registrar
   MAY be administratively configured to support only a particular DTLS
   version or higher.

   An EST-coaps server [RFC9148] (if present as a separate entity from
   above Registrar) that implements this specification also MUST support
   both DTLS 1.3 and DTLS 1.2 client connections by default.  However,
   for security reasons the EST-coaps server MAY be administratively
   configured to support only a particular DTLS version or higher.

6.1.2.  TLS Client Certificates: IDevID authentication

   As described in Section 5.1 of [RFC8995], the Pledge makes a
   connection to the Registrar using a TLS Client Certificate for
   authentication.  This is the Pledge's IDevID certificate.

   Subsequently the Pledge will send a Pledge Voucher Request (PVR).
   Further elements of Pledge authentication may be present in the PVR,
   as detailed in Section 9.2.






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6.1.3.  DTLS Handshake Fragmentation Considerations

   DTLS includes a mechanism to fragment handshake messages.  This is
   described in Section 4.4 of [RFC9147]. cBRSKI will often be used with
   a Join Proxy, described in Section 6.2, which relays each DTLS
   message to the Registrar.  A stateless Join Proxy will need some
   additional space to wrap each DTLS message inside a Join Proxy UDP
   message, while the wrapped result needs to fit in the maximum IPv6
   MTU guaranteed on 6LoWPAN [RFC6282] networks, which is 1280 bytes.

   For this reason it is RECOMMENDED that a PMTU of 1024 bytes be
   assumed for the DTLS handshake and appropriate DTLS fragmentation is
   used.  It is unlikely that any ICMPv6 Packet Too Big indications
   ([RFC4443]) will be relayed by the Join Proxy back to the Pledge.

   During the operation of the EST-coaps protocol, the CoAP Block-wise
   transfer mechanism [RFC7959] will be automatically used when message
   sizes exceed the PMTU.  A Pledge/EST-client on a constrained network
   MUST use the (D)TLS maximum fragment length extension
   ('max_fragment_length') defined in Section 4 of [RFC6066] with the
   maximum fragment length set to a value of either 2^9 or 2^10, when
   operating as a DTLS 1.2 client.

   A Pledge/EST-client operating as DTLS 1.3 client, MUST use the (D)TLS
   record size limit extensions ('record_size_limit') defined in
   Section 4 of [RFC8449], with RecordSizeLimit set to a value between
   512 and 1024 (inclusive).

6.1.4.  Registrar and the Server Name Indicator (SNI)

   The SNI issue described below affects [RFC8995] as well, and is
   reported in errata: https://www.rfc-editor.org/errata/eid6648
   (https://www.rfc-editor.org/errata/eid6648)

   As the Registrar is discovered by IP address, and typically connected
   via a Join Proxy, the hostname of the Registrar is not known to the
   Pledge.  Therefore, it cannot do DNS-ID validation ([RFC9525]) on the
   Registrar's certificate.  Instead, it must do validation using the
   voucher.

   Without knowing the hostname, the Pledge cannot put any reasonable
   value into the [RFC6066] Server Name Indicator (SNI) extension.
   Therefore the Pledge SHOULD omit the SNI extension as per Section 9.2
   of [RFC8446].

   In some cases, particularly while testing BRSKI, a Pledge may be
   given the hostname of a particular Registrar to connect to directly.
   Such a bypass of the discovery process may result in the Pledge



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   taking a different code branch to establish a DTLS connection, and
   may result in the SNI being inserted by a library.  For this reason,
   the Registrar MUST ignore any SNI it receives from a Pledge.

   A primary motivation for making the SNI ubiquitous in the public web
   is because it allows for multi-tenant hosting of HTTPS sites on a
   single (scarce) IPv4 address.  This consideration does not apply to
   the server function in the Registrar because:

   *  it uses DTLS and CoAP, not HTTPS;

   *  it typically uses IPv6, often [RFC4193] Unique Local Address,
      which are plentiful;

   *  the server port number is typically discovered, so multiple
      tenants can be accommodated via unique UDP port numbers.

6.1.5.  Registrar Server Certificate Requirements

   As per Section 3.6.1 of [RFC7030], the Registrar certificate MUST
   have the Extended Key Usage (EKU) id-kp-cmcRA.  This certificate is
   also used as a TLS Server Certificate, so it MUST also have the EKU
   id-kp-serverAuth.

   See Appendix C.2.2 for an example of a Registrar certificate with
   these EKUs set.  See Section 7.4 for Registrar client certificate
   requirements.

6.2.  cBRSKI Join Proxy

   [I-D.ietf-anima-constrained-join-proxy] specifies the details for a
   stateful or stateless constrained Join Proxy which is equivalent to
   the BRSKI Proxy defined in [RFC8995], Section 4.  See also Section 10
   for more details on discovery of a Join Proxy by a Pledge, and
   discovery of a Registrar by a Join Proxy.

6.3.  Request URIs, Resource Discovery and Content-Formats

   cBRSKI operates using CoAP over DTLS, with request URIs using the
   coaps scheme.  The Pledge operates in CoAP client role.  To keep the
   protocol messages small the EST-coaps and cBRSKI request URIs are
   shorter than the respective EST and BRSKI URIs.

   During the cBRSKI onboarding on an IPv6 network these request URIs
   have the following form:

     coaps://[<link-local-ipv6>]:<port>/.well-known/brski/<short-name>
     coaps://[<link-local-ipv6>]:<port>/.well-known/est/<short-name>



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   where <link-local-ipv6> is the discovered link-local IPv6 address of
   a Join Proxy, and <port> is the discovered port of the Join Proxy
   that is used to offer the cBRSKI proxy functionality.

   <short-name> is the short resource name for the cBRSKI and EST-coaps
   resources.  For EST-coaps, Section 5.1 of [RFC9148] defines the CoAP
   <short-name> resource names.  For cBRSKI, this document defines the
   short resource names based on the [RFC8995] long HTTP resource names.
   See Table 1 for a summary of these resource names.

   Section 11 details how the Pledge discovers a Join Proxy link-local
   address and port in different deployment scenarios.

   The request URI formats defined here enable the Pledge to perform
   onboarding/enrollment without requiring discovery of the available
   onboarding options, voucher formats, BRSKI/EST resources, enrollment
   protocols, and so on.  This is helpful for the majority of
   constrained Pledges that would support only a single set of these
   options.  However, for Pledges that do support multiple options,
   [I-D.ietf-anima-brski-discovery] will define discovery methods so
   that a Pledge can select the optimal set of options for the current
   onboarding operation.

   Alternatively, a Pledge could also send CoAP discovery queries
   (Section 7 of [RFC7252]) to the Registrar to discover detailed
   options for onboarding and/or enrollment functions.  Supporting these
   queries is OPTIONAL for both the Pledge and the Registrar.  To
   clarify which options in particular can be discovered, Appendix E
   provides an informative overview of what can be discovered and how to
   discover it.

   Because a Pledge only has indirect access to the Registrar via a
   single port on the Join Proxy, the Registrar MUST host all cBRSKI/
   EST-coaps resources on the same (UDP) server IP address and port.
   This is the address and port where a Join Proxy would relay DTLS
   records from the Pledge to.

   Although the request URI templates include IP address, scheme and
   port, in practice the CoAP request message sent over the secure DTLS
   connection only encodes the URI path explicitly.  For example, a
   Pledge that skips resource discovery operations just sends the
   initial CoAP voucher request as follows:

     REQ: POST /.well-known/brski/rv
       Content-Format: 836 (application/voucher+cose)
       Payload       : (COSE-signed Pledge Voucher Request, PVR)





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   Note that only content-format 836 (application/voucher+cose) is
   defined in this document for the payload sent to the voucher request
   resource (/rv).  Content-format 836 MUST be supported by the
   Registrar for the /rv resource and it MAY support additional formats.
   The Pledge MAY also indicate in the request the desired format of the
   (voucher) response, using the Accept Option.  An example of using
   this option in the request is as follows:

     REQ: POST /.well-known/brski/rv
       Content-Format: 836 (application/voucher+cose)
       Accept        : 836 (application/voucher+cose)
       Payload       : (COSE-signed Pledge Voucher Request, PVR)

   If the Accept Option is omitted in the request, the response format
   follows from the request payload format (which is 836).

   Note that this specification allows for application/voucher+cose
   format requests and vouchers to be transported over HTTPS, as well as
   for application/voucher-cms+json and other formats yet to be defined
   over CoAP.  The burden for this flexibility is placed upon the
   Registrar.  A Pledge on constrained hardware is expected to support a
   single format only.

   The Pledge and MASA need to support one or more formats (at least
   format 836) for the voucher and for the voucher request.  The MASA
   needs to support all formats that the Pledge supports.

6.3.1.  Status Telemetry Returns

   [RFC8995] defines two telemetry returns from the Pledge which are
   sent to the Registrar.  These are the BRSKI Status Telemetry
   [RFC8995], Section 5.7 and the Enrollment Status Telemetry [RFC8995],
   Section 5.9.4.  These are two CoAP POST requests made the by Pledge
   at two key steps in the process.

   [RFC8995] defines the content of these POST operations in CDDL, which
   are serialized as JSON.  This document extends this with an
   additional CBOR format, derived using the CDDL rules in [RFC8610].

   The new CBOR telemetry format has CoAP content-format 60
   (application/cbor) and MUST be supported by the Registrar for both
   the /vs and /es resources.  The existing JSON format has CoAP
   content-format 50 (application/json) and MAY also be supported by the
   Registrar.  A Pledge MUST use the new CBOR format to send telemetry
   messages.






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6.3.2.  CoAP Resources Table

   cBRSKI inherits EST-coaps [RFC9148] functions: specifically, the
   mandatory Simple (Re-)Enrollment (/sen and /sren) and Certification
   Authority certificates request (/crts).  Support for CSR Attributes
   Request (/att) and server-side key generation (/skg, /skc) remains
   optional for the EST-coaps server.

   Table 1 summarizes the resources used in cBRSKI.  It includes both
   the short-name cBRSKI resources and the EST-coaps resources.

   +=================+====================+===============+============+
   | BRSKI + EST     | cBRSKI + EST-coaps | Well-known    | Required   |
   | name            | <short-name>       | URI           | for        |
   |                 |                    | namespace     | Registrar? |
   +=================+====================+===============+============+
   | /enrollstatus   | /es                | brski         | MUST       |
   +-----------------+--------------------+---------------+------------+
   | /requestvoucher | /rv                | brski         | MUST       |
   +-----------------+--------------------+---------------+------------+
   | /voucher_status | /vs                | brski         | MUST       |
   +-----------------+--------------------+---------------+------------+
   | /cacerts        | /crts              | est           | MUST       |
   +-----------------+--------------------+---------------+------------+
   | /csrattrs       | /att               | est           | MAY        |
   +-----------------+--------------------+---------------+------------+
   | /simpleenroll   | /sen               | est           | MUST       |
   +-----------------+--------------------+---------------+------------+
   | /simplereenroll | /sren              | est           | MUST       |
   +-----------------+--------------------+---------------+------------+
   | /serverkeygen   | /skg               | est           | MAY        |
   +-----------------+--------------------+---------------+------------+
   | /serverkeygen   | /skc               | est           | MAY        |
   +-----------------+--------------------+---------------+------------+

        Table 1: BRSKI/EST resource name mapping to cBRSKI/EST-coaps
                            short resource name

6.3.3.  CoAP Uri-Path Abbreviation

   To minimize the size of CoAP request packets on constrained networks,
   the CoAP Uri-Path-Abbrev Option defined in
   [I-D.ietf-core-uri-path-abbrev] MUST be supported by the Registrar.

6.4.  CoAP Responses

   [RFC8995], Section 5 defines a number of HTTP response codes that the
   Registrar is to return when certain conditions occur.



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   The 401, 403, 404, 406 and 415 response codes map directly to CoAP
   codes 4.01, 4.03, 4.04, 4.06 and 4.15 respectively.

   The 202 Retry process which may occur in the voucher request, is to
   be handled in the same way as the Section 5.7 of [RFC9148] process
   for Delayed Responses.

6.5.  Extensions to EST-coaps

   This section defines extensions to EST-coaps for Pledges (during
   initial onboarding), EST-coaps clients (after initial onboarding) and
   Registrars (that implement an EST-coaps server).  Note that a device
   that is already onboarded is not called "Pledge" in this section: it
   now acts in the role of an EST-coaps client.

6.5.1.  Pledge Enrollment Procedure

   This section defines optimizations for the EST-coaps protocol as used
   by a Pledge.  These aim to reduce payload sizes and the number of
   messages (round-trips) required for the initial EST enrollment.

   A Pledge SHOULD NOT perform the optional EST-coaps "CSR attributes
   request" (/att).  Instead, the Pledge selects the attributes to
   include in the CSR as specified below.

   One or more Subject Distinguished Name fields MUST be included in the
   CSR.  If the Pledge has no specific information on what attributes/
   fields are desired in the CSR, which is the common case, it MUST use
   the Subject Distinguished Name fields from its IDevID unmodified.
   Note that a Pledge MAY receive such specific information via the
   voucher data (encoded in a vendor-specific way, or as defined by a
   future specification) or via some other, out-of-band means.

   A Pledge uses the following optimized EST-coaps procedure:

   1.  If the voucher, that validates the current Registrar, contains a
       single pinned domain CA certificate, the Pledge provisionally
       considers this certificate as the EST trust anchor, as if it were
       the result of a "CA certificates request" (/crts) to the
       Registrar.

   2.  Using this CA certificate as trust anchor it proceeds with EST
       simple enrollment (/sen) to obtain a provisionally trusted LDevID
       certificate.

   3.  If the Pledge determines that the pinned domain CA is (1) a root
       CA certificate and (2) signer of the LDevID certificate, the
       Pledge accepts the pinned domain CA certificate as the legitimate



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       trust anchor root CA for the Registrar's domain.  It also accepts
       the LDevID certificate as its new LDevID identity.  And steps 4
       and 5 are skipped.

   4.  Otherwise, if the step 3 condition was not met, the Pledge MUST
       perform a "CA certificates request" (/crts) to the EST server to
       obtain the full set of EST CA trust anchors.  It then MUST
       attempt to chain the LDevID certificate to one of the CAs in the
       set.

   5.  If the Pledge cannot obtain the set of CA certificates, or it is
       unable to create the chain as defined in step 4, the Pledge MUST
       abort the enrollment process and report the error using the
       enrollment status telemetry (/es).

6.5.2.  Renewal of CA certificates

   An EST-coaps client that has an idea of the current time (internally,
   or via Network Time Protocol) SHOULD consider the validity time of
   the trust anchor CA(s), and MAY begin requesting new trust anchor
   certificates(s) using the /crts request when the CA has 50% of it's
   validity time (notAfter - notBefore) left.  A client without access
   to the current time cannot decide if trust anchor CA(s) have expired,
   and SHOULD poll periodically for a new trust anchor certificate(s)
   using the /crts request at an interval of approximately 1 month.  An
   EST-coaps server SHOULD include the CoAP ETag Option ([RFC7252],
   Section 5.10.6)in every response to a /crts request, to enable
   clients to perform low-overhead validation whether their trust anchor
   CA is still valid.  The EST-coaps client SHOULD store the ETag
   resulting from a /crts response in memory and SHOULD use this value
   in an ETag Option in its next GET /crts request.

6.5.3.  Change of Domain Trust Anchor(s)

   The domain trust anchor(s) may change over time.  Such a change may
   happen due to relocation of the client device to a new domain, a new
   subdomain, or due to a key update of a trust anchor as described in
   [RFC4210], Section 4.4.

   From the client's viewpoint, a trust anchor change happens during
   EST-coaps re-enrollment: since a change of domain CA requires all
   devices operating under the old domain CA to acquire a new LDevID
   certificate issued by the new domain CA.  A client's re-enrollment
   may be triggered by various events, such as an instruction to re-
   enroll sent by a domain entity, or an imminent expiry of its LDevID
   certificate, or other.  How the re-enrollment is explicitly triggered
   on the client by a domain entity, such as a commissioner or a
   Registrar, is out of scope of this specification.



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   The mechanism described in [RFC7030], Section 4.1.3 and [RFC4210],
   Section 4.4 for root CA key update requires four certificates:
   OldWithOld, OldWithNew, NewWithOld, and NewWithNew.  Of these four,
   the OldWithOld certificate is already stored in the client's Explicit
   TA database.  The other certificates will be provided to the client
   in a /crts response, during the EST-coaps re-enrollment procedure.

6.5.4.  Re-enrollment Procedure

   For re-enrollment, the EST-coaps client MUST support the following
   EST-coaps procedure.  During this procedure the EST-coaps server MAY
   re-enroll the client into a new domain or into a new sub-CA within a
   larger domain.

   1.  The client connects with DTLS to the EST-coaps server, and
       authenticates with its present domain certificate (LDevID) as
       usual.  The EST-coaps server authenticates itself with its domain
       RA certificate that is currently trusted by the client, i.e. it
       chains to a trust anchor CA that the client has stored in its
       Explicit TA database.  This is the OldWithOld trust anchor.  The
       client checks that the server is a Registration Authority (RA) of
       the domain as required by Section 3.6.1 of [RFC7030] before
       proceeding.

   2.  The client performs the simple re-enrollment request (/sren) and
       upon success it obtains a new LDevID certificate.

   3.  The client verifies the new LDevID certificate against its
       Explicit TA database.  If the new LDevID chains successfully to a
       TA, this means trust anchors did not significantly change and the
       client MAY skip retrieving the current CA certificates using the
       "CA certificates request" (/crts).  If it does not chain
       successfully, it means trust anchor(s) were changed significantly
       and the client MUST retrieve the new domain trust anchors using
       the "CA certificates request" (/crts).

   4.  If the client retrieved new trust anchor(s) in step 3, then it
       MUST verify that the new LDevID certificate it obtained in step 2
       chains with the new trust anchor(s).  If it chains successfully,
       the client stores the new trust anchor(s) in its Explicit TA
       database, accepts the new LDevID certificate and stops using its
       prior LDevID certificate.  If it does not chain successfully, the
       client MUST NOT update its LDevID certificate, and it MUST NOT
       update its Explicit TA database, and the client MUST abort the
       enrollment process and MUST attempt to report the error to the
       EST-coaps server using enrollment status telemetry (/es).





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   Note that even though the EST-coaps client may skip the /crts request
   in step 3 at this time, it SHOULD still support renewal of the trust
   anchors as detailed in Section 6.5.2.

   Note that an EST-coaps server that is also a Registrar will already
   support the enrollment status telemetry resource (/es) in step 4,
   while an EST-coaps server that purely implements [RFC9148], and not
   the present specification, will not support this resource.

6.5.5.  Multipart Content-Format for CA certificates (/crts) Resource

   In EST-coaps [RFC9148] the PKCS#7 container format is used for CA
   certificates distribution.  Because the PKCS#7 format is only used as
   a certificate container and no additional security is applied on the
   container, it becomes attractive to replace this format by something
   simpler, on a constrained Pledge: so that additional PKCS#7 code is
   avoided.  Therefore, this document defines a container format using
   the [RFC8710] application/multipart-core media type (CoAP content-
   format 62).  This is beneficial since a Pledge necessarily already
   supports CBOR parsing, so there is little code overhead to support
   this CBOR-based container format.

   A Registrar or EST-coaps server MUST support content-format 62 for
   the /crts resource.  The multipart collection MUST contain the
   individual CA certificates, each encoded as an application/pkix-cert
   (287) representation.  Future documents may define other certificate
   formats: the multipart collection can handle any future types.  The
   order of CA certificates MUST be in the CA hierarchy order starting
   from the issuer of the client's LDevID first, up to the highest-level
   domain CA, then optionally followed by any further CA certificates
   that are not part of this hierarchy.  These further CA certificates
   may be Third-party TAs as defined in [RFC7030].  The highest-level
   domain CA may or may not be a root CA certificate.

   As an example, for the two-level CA domain PKI of Figure 1 the
   multipart container will contain two representations:

   [ <domain sub-CA cert (2)> , <domain CA cert (1)> ]

   Encoded as an application/multipart-core CBOR array this is (shown in
   CBOR diagnostic notation):

   [ 287, h'3082' ... 'd713', 287, h'3082' ... 'a034' ]

   The total number of CA certificates SHOULD be 1, 2 or 3 and not
   higher to prevent constrained Pledges from running out of memory for
   the trust anchor storage (Explicit TA database).  However if a domain
   operator can guarantee that any Pledges enrolled in its network can



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   support larger sets of CA certificates, the total number MAY be
   configured as higher than 3.  To facilitate a reliable transfer of
   large payloads over constrained networks, the server MUST support
   CoAP Block-wise transfer for the /crts response.  The server MUST
   also support the Size2 Option [RFC7959] to provide the total resource
   length in bytes, when requested by a client.

   Implementation notes: a client that receives the first block of
   payload data from the server, can already inspect the total number of
   CA certificates by decoding the first byte of the payload.  In CBOR
   encoding, the respective first bytes 0x81-0x97 represent an array
   with length 1-23, respectively.  Furthermore, the length in bytes of
   the first CA certificate can be already determined by decoding the
   first bytes of the second element, because the CBOR encoding for
   binary string includes the length of this string.  A client that
   requires an estimate of the total resource size (to be returned as
   part of the first Block2 response from the server) can use a Size2
   Option with value 0 in its request.  Knowing the overall progress of
   the data transfer may be helpful in certain cases, e.g. when a Pledge
   provides visual progress information on the onboarding progress.

6.6.  Registrar Extensions

   Before a Registrar forwards a COSE-signed voucher from MASA to the
   Pledge, it MUST remove any 'x5bag' or 'x5chain' unprotected COSE
   header attributes (which are defined in [RFC9360]).  The contents of
   these unprotected attributes are solely for validation/logging use by
   the Registrar.  Removing these attributes reduces the voucher size on
   the constrained network path to the Pledge.

   The content-format 60 (application/cbor) MUST be supported by the
   Registrar for the /vs and /es resources.

   Content-format 836 (application/voucher+cose) MUST be supported by
   the Registrar for the /rv resource for CoAP POST requests, both as
   request payload and as response payload.

   Content-format 287 (application/pkix-cert) MUST be supported by the
   Registrar as a response payload for the /sen and /sren resources.












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   When a Registrar receives a "CA certificates request" (/crts) request
   with a CoAP Accept Option with value 287 (application/pkix-cert) it
   MUST return only the single CA certificate that is the envisioned or
   actual CA authority for the current, authenticated Pledge making the
   request.  An exception to this rule is when the domain has been
   configured to operate with multiple CA trust anchors exclusively:
   then the Registrar returns a 4.06 Not Acceptable error to signal to
   the client that it needs to request another content-format that
   supports retrieval of multiple CA certificates.

7.  BRSKI-MASA Protocol

   This section describes extensions to and clarifications of the BRSKI-
   MASA protocol between Registrar and MASA.

7.1.  Protocol and Formats

   Section 5.4 of [RFC8995] describes a connection between the Registrar
   and the MASA as being a normal TLS connection using HTTPS.  This
   document does not change that.

   The MASA only needs to support formats for which it has constructed
   Pledges that use that format.

   The Registrar MUST use the same format for the RVR as the Pledge used
   for its PVR.  Specifically, the Registrar MUST use the media type
   application/voucher+cose for its voucher request to MASA, when the
   Pledge used content-format 836 in the payload of its request to the
   Registrar.

   The Registrar indicates the voucher format (by media type) it wants
   to receive from MASA using the HTTP Accept header.  This format MUST
   be the same as the format of the PVR, so that the Pledge can parse
   the resulting voucher.

   At the moment of writing the creation of CoAPS based MASAs is deemed
   unrealistic and unnecessary.  The use of CoAP for the BRSKI-MASA
   connection is out of scope but can be the subject of another
   document.  Some consideration was made to specify CoAP support for
   consistency, but:

   *  the Registrar is not expected to be so constrained that it cannot
      support HTTPS client connections.

   *  the technology and experience to build Internet-scale HTTPS
      responders (which the MASA is) is common, while the experience
      doing the same for CoAP is much less common.




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   *  a Registrar is likely to provide onboarding services to both
      constrained and non-constrained devices.  Such a Registrar would
      need to speak HTTPS anyway.

   *  a manufacturer is likely to offer both constrained and non-
      constrained devices, so there may in practice be no situation in
      which the MASA could be CoAP-only.  Additionally, as the MASA is
      intended to be a function that can easily be outsourced to a
      third-party service provider, reducing the complexity would also
      seem to reduce the cost of that function.

   *  security-related considerations: see Section 14.3.

7.2.  Registrar Voucher Request

   If the PVR contains a proximity assertion, the Registrar MUST
   propagate this assertion into the RVR by including the 'assertion'
   attribute with the value "proximity".  This conforms to the example
   in Section 3.3 of [RFC8995] of carrying the assertion forward.

7.3.  MASA and the Server Name Indicator (SNI)

   A TLS/HTTPS connection is established between the Registrar and MASA.

   Section 5.4 of [RFC8995] explains this process, and there are no
   externally visible changes made by this document.  A MASA that
   supports the unconstrained voucher formats should be able to support
   constrained voucher formats equally well.

   There is no requirement that a single MASA be used for both
   constrained and unconstrained voucher requests: the choice of MASA is
   determined by the id-mod-MASAURLExtn2016 extension contained in the
   IDevID, so it can be determined by the manufacturer.

   The Registrar MUST do DNS-ID checks ([RFC9525]) on the contents of
   the certificate provided by the MASA during the TLS handshake.

   In contrast to the Pledge/Registrar situation, the Registrar always
   knows the name of the MASA, and MUST always include an [RFC6066]
   Server Name Indicator.  The SNI is optional in TLS 1.2, but common.
   The SNI is considered mandatory with TLS 1.3.

   The presence of the SNI extension is required by the MASA, in order
   for the MASA's server to host multiple tenants (for different
   customers).






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7.4.  Registrar Client Certificate Requirements

   The Registrar SHOULD use a TLS Client Certificate to authenticate to
   the MASA per Section 5.4.1 of [RFC8995].  If the certificate that the
   Registrar uses is marked as a id-kp-cmcRA certificate, via Extended
   Key Usage, then it MUST also have the id-kp-clientAuth EKU attribute
   set.

   In summary, for typical Registrar use, where a single Registrar
   certificate is used for both client and server roles, the certificate
   MUST have an EKU set with at least all of id-kp-cmcRA, id-kp-
   serverAuth, and id-kp-clientAuth.

8.  Pinning in Voucher Artifacts

   The voucher is a statement by the MASA for use by the Pledge that
   provides the identity of the Pledge's owner.  This section describes
   how the owner's identity is determined and how it is specified within
   the voucher.

8.1.  Registrar Identity Selection and Encoding

   Section 5.5 of [RFC8995] describes BRSKI policies for selection of
   the owner identity.  It indicates some of the flexibility enabled for
   the Registrar, and recommends the Registrar to include only
   certificates in the voucher request (CMS) signing structure that
   participate in the certificate chain that is to be pinned.

   The MASA is expected to evaluate the certificates included in an RVR,
   forming them into a chain with the Registrar's (signing) identity on
   one end.  Then, it pins a certificate selected from this chain
   according to its pinning policy (Section 8.2).

   For instance, for a domain with a two-level certification authority
   (see Figure 1), where the RVR has been signed by "domain Registrar",
   the RVR includes the chain formed by the domain Registrar EE
   certificate, the domain Sub-CA certificate, and the domain CA trust
   anchor certificate.  The arrows in the figure indicate the issuing of
   a certificate, i.e. author of (1) issued (2) and author of (2) issued
   (3).











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    .------------------.
    |  domain CA (1)   |
    |  trust anchor    |
    '------------------'
              |
              v
       .------------.
       | domain (2) |
       | Sub-CA     |
       '------------'
              |
              v
     .----------------.
     |   domain       |
     | Registrar (3)  |
     | EE certificate |
     '----------------'

                         Figure 1: Two-Level CA PKI

   When the Registrar is using a COSE-signed RVR, the COSE_Sign1 object
   contains a protected and an unprotected header.  The Registrar MUST
   place all the certificates needed by MASA to validate the signature
   chain for its RVR in an 'x5bag' attribute in either the protected or
   the unprotected header as defined in Section 2 of [RFC9360].

8.2.  MASA Pinning Policy

   The MASA, having assembled and verified the certificate chain that
   signed the RVR then needs to select a certificate to pin.  (For the
   case that only the Registrar's End-Entity certificate is included,
   only this certificate can be selected and this section does not
   apply.)  The BRSKI policy for pinning by the MASA as described in
   Section 5.5.2 of [RFC8995] leaves much flexibility to the
   manufacturer.

   The present document adds the following rules to the MASA pinning
   policy to reduce the network load on the constrained network side:

   1.  for a voucher containing a nonce, it SHOULD pin the most specific
       (lowest-level) CA certificate in the chain.

   2.  for a nonceless voucher, it SHOULD pin the least-specific
       (highest-level) CA certificate in the chain that is allowed under
       the MASA's policy for this specific domain.






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   The rationale for 1. is that in case of a voucher with nonce, the
   voucher is valid only in scope of the present DTLS connection between
   Pledge and Registrar anyway, so there is no benefit to pin a higher-
   level CA.  By pinning the most specific CA the constrained Pledge can
   validate its DTLS connection using less crypto operations.  The
   rationale for pinning a CA instead of the Registrar's End-Entity
   certificate directly is based on the following benefit on constrained
   networks: the pinned certificate in the voucher can in common cases
   be re-used as a Domain CA trust anchor during the EST enrollment and
   during the operational phase that follows after EST enrollment, as
   explained in Section 6.5.1.

   The rationale for 2. follows from the flexible BRSKI trust model for,
   and purpose of, nonceless vouchers (Sections 5.5.* and 7.4.1 of
   [RFC8995]).

   Referring to the example of Figure 1 of a domain with a two-level
   certification authority, the most specific CA ("Sub-CA") is the
   identity that is pinned by MASA in a nonced voucher.

   In case of a nonceless voucher, depending on the trust level, the
   MASA pins the "Registrar" certificate (low trust in customer), or the
   "Sub-CA" certificate (in case of medium trust, implying that any
   Registrar of that sub-domain is acceptable), or even the "domain CA"
   certificate (in case of high trust in the customer, and possibly a
   pre-agreed need of the customer to obtain flexible long-lived
   vouchers).

8.3.  Pinning of Raw Public Keys (RPK)

   Specifically for the most-constrained use cases, the pinning of the
   raw public key (RPK) of the Registrar is also supported in the
   constrained voucher, instead of a PKIX certificate.  This is used by
   the RPK variant of cBRSKI described in Section 13, but it can also be
   used in the default cBRSKI flow as a means to reduce voucher size.

   For both cases, if an RPK is pinned, it MUST be the RPK of the
   Registrar, which equals the public key of the Registrar's EE
   certificate.

   When the Pledge is known by MASA to support the RPK variant only, the
   voucher produced by the MASA pins the RPK of the Registrar in either
   the 'pinned-domain-pubk' or 'pinned-domain-pubk-sha256' attribute of
   the voucher data.  This is described in more detail in [RFC8366bis]
   and Section 13.






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   When the Pledge is known by MASA to support PKIX operations, the
   'pinned-domain-cert' attribute present in a voucher normally pins a
   domain certificate.  That can be either the End-Entity certificate of
   the Registrar, or the certificate of a domain CA, as specified in
   Section 8.2.  However, if the Pledge is known by MASA to also support
   RPK pinning and the MASA policy intends to pin the Registrar in the
   voucher (and not a CA), then MASA SHOULD pin the RPK (RPK3 in
   Figure 2) of the Registrar instead of the Registrar's End-Entity
   certificate to save space in the voucher.

                                             .-------------.
    .------------.                           | private     |
    | pub-CA (1) |                           | root-CA (1) |
    '------------'                           '-------------'
           |                                        |
           v             .-------------.            v
    .------------.       | private     |     .------------.
    | sub-CA (2) |       | root-CA (1) |     | sub-CA (2) |
    '------------'       '-------------'     '------------'
           |                    |                   |
           v                    v                   v
   .--------------.     .--------------.     .--------------.
   | Registrar(3) |     | Registrar(3) |     | Registrar(3) |
   |    RPK3      |     |    RPK3      |     |    RPK3      |
   '--------------'     '--------------'     '--------------'

              Figure 2: Raw Public Key (RPK) pinning examples

9.  Artifacts

   The YANG ([RFC7950]) module and CBOR serialization for the
   constrained voucher as used by cBRSKI are described in [RFC8366bis].
   That document also assigns SID values to YANG elements in accordance
   with [RFC9254] and [RFC9595].  The present section provides some
   examples of these artifacts and defines a new signature format for
   these, using COSE.

   Compared to the first voucher request definition done in [RFC8995],
   the constrained voucher request adds the attributes 'proximity-
   registrar-pubk' and 'proximity-registrar-pubk-sha256'.  One of these
   is optionally used to replace the 'proximity-registrar-cert'
   attribute, for a smaller voucher request data size - useful for the
   most constrained cases.

   The constrained voucher adds the attributes 'pinned-domain-pubk' and
   'pinned-domain-pubk-sha256' to pin an RPK.  One of these is
   optionally used instead of the 'pinned-domain-cert' attribute, for a
   smaller voucher data size.



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9.1.  Example Artifacts

9.1.1.  Example Pledge Voucher Request (PVR) Artifact

   Below, example voucher data from a constrained voucher request (PVR)
   from a Pledge to a Registrar is shown in CBOR diagnostic notation.
   Long CBOR byte strings have been shortened for readability, using the
   ellipsis ("...") to indicate elided bytes.  This notation is defined
   in [I-D.ietf-cbor-edn-literals].  The enum value of the assertion
   attribute is 2 for the 'proximity' assertion as defined in
   Section 8.3 of [RFC8366bis].

   {
    2501: {          / SID=2501, ietf-voucher-request:voucher|voucher /
      1: 2,                      / SID=2502, assertion 2 = "proximity"/
      7: h'831D5198A6CA2C7F',    / SID=2508, nonce                    /
     12: h'30593013' ... '9A54', / SID=2513, proximity-registrar-pubk /
     13: "JADA123456789"         / SID=2514, serial-number            /
    }
   }

   The Pledge has included the attribute 'proximity-registrar-pubk'
   which carries the public key of the Registrar, instead of including
   the full Registrar certificate in a 'proximity-registrar-cert'
   attribute.  This is done to reduce the size of the PVR.  Also note
   that the Pledge did not include the 'created-on' attribute since it
   lacks an internal real-time clock and has no knowledge of the current
   time at the moment of performing the onboarding.

9.1.2.  Example Registrar Voucher Request (RVR) Artifact

   Next, example voucher data from a constrained voucher request (RVR)
   from a Registrar to a MASA is shown in CBOR diagnostic notation.  The
   Registrar has created this request triggered by the reception of the
   Pledge voucher request (PVR) of the previous example.  Again, long
   CBOR byte strings have been shortened for readability.

   {
    "ietf-voucher-request:voucher": {
       "assertion":     2,
       "created-on":    "2022-12-05T19:19:19.536Z",
       "nonce":         h'831D5198A6CA2C7F',
       "idevid-issuer": h'04183016' ... '1736C3E0',
       "serial-number": "JADA123456789",
       "prior-signed-voucher-request": h'A11909' ... '373839'
    }
   }




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   Note that the Registrar uses here the string data type for all key
   names, instead of the more compact SID integer keys.  This is fine
   for any use cases where the network between Registrar and MASA is an
   unconstrained network where data size is not critical.  The
   constrained voucher request format supports both the string and SID
   key types, for PVR as well as RVR.

9.1.3.  Example Voucher Artifacts

   Below, an example of constrained voucher data is shown in CBOR
   diagnostic notation.  It was created by a MASA in response to
   receiving the Registrar Voucher Request (RVR) shown in Section 9.1.2.
   The enum value of the 'assertion' attribute is set to "proximity"
   (2), to acknowledge to both the Pledge and the Registrar that the
   proximity of the Pledge to the Registrar is considered proven.

   {
    2451: {               / SID = 2451, ietf-voucher:voucher|voucher /
      1: 2,                      / SID = 2452, assertion "proximity" /
      2: "2022-12-05T19:19:23Z", / SID = 2453, created-on            /
      3: false,       / SID = 2454, domain-cert-revocation-checks    /
      7: h'831D5198A6CA2C7F',    / SID = 2508, nonce                 /
      8: h'308201' ... '8CFF',   / SID = 2459, pinned-domain-cert    /
     11: "JADA123456789"         / SID = 2462, serial-number         /
    }
   }

   While the above example voucher data includes the nonce from the PVR,
   the next example is for a nonce-less voucher.  Instead of a nonce, it
   includes an 'expires-on' attribute with the date and time on which
   the voucher expires.  Because the MASA did not verify the proximity
   of the Pledge and Registrar in this case, the 'assertion' attribute
   contains a weaker assertion of "verified" (0).  This indicates that
   the MASA verified the domain's ownership of the Pledge via some other
   means.

   {
    2451: {               / SID = 2451, ietf-voucher:voucher|voucher /
      1: 0,                      / SID = 2452, assertion "verified"  /
      2: "2022-12-06T10:15:32Z", / SID = 2453, created-on            /
      3: false,          / SID = 2454, domain-cert-revocation-checks /
      4: "2022-12-13T10:15:32Z", / SID = 2455, expires-on            /
      8: h'308201F8' ... 'FF',   / SID = 2459, pinned-domain-cert    /
     11: "JADA123456789"         / SID = 2462, serial-number         /
    }
   }





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   The voucher is valid for one week.  To verify the voucher's validity,
   the Pledge would either need an internal real-time clock or some
   external means of obtaining the current time, such as Network Time
   Protocol (NTP) or a radio time signal receiver.

9.2.  Signing Voucher and Voucher Request Artifacts with COSE

   The COSE_Sign1 structure is discussed in Section 4.2 of [RFC9052].
   The CBOR object that carries the body, the signature, and the
   information about the body and signature is called the COSE_Sign1
   structure.  It is used when only one signature is used on the body.

   Support for ECDSA with SHA2-256 using curve secp256r1 (aka
   prime256k1) is RECOMMENDED.  Most current low power hardware has
   support for acceleration of this algorithm.  Future hardware designs
   could omit this in favour of a newer algorithms.  This is the ES256
   (-7) keytype from Table 1 of [RFC9053].  Support for curve secp256k1
   is OPTIONAL.

   Support for EdDSA using Curve 25519 is RECOMMENDED in new designs if
   hardware support is available.  This is keytype EDDSA (-8) from
   Table 2 of [RFC9053].  A 'crv' parameter is necessary to specify the
   Curve, for example value Ed25519 (6) from Table 18 of [RFC9053].  The
   'kty' field MUST be present, and it MUST be "OKP" (Table 17 of
   [RFC9053]).

   A transition towards EdDSA is occurring in the industry.  Some
   hardware can accelerate only some algorithms with specific curves,
   other hardware can accelerate any curve, and still other kinds of
   hardware provide a tool kit for acceleration of any elliptic curve
   algorithm.

   In general, the Pledge is expected to support only a single
   algorithm, while the Registrar, usually not constrained, is expected
   to support a wide variety of algorithms: both legacy ones and up-and-
   coming ones via regular software updates.

   An example of the supported COSE_Sign1 object structure containing a
   Pledge Voucher Request (PVR) is shown below.












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   18(                   / tag for COSE_Sign1                       /
    [
      h'A10126',         / protected COSE header encoding: {1: -7}  /
                         /            which means { "alg": ES256 }  /
      {},                / unprotected COSE header parameters       /
      h'A119' ... '3839', / PVR payload wrapped in CBOR byte string /
      h'4567' ... '1234'  / PVR binary Sign1 signature              /
    ]
   )

   The [COSE-registry] specifies the integers/encoding for the 'alg'
   field.  The 'alg' field restricts the key usage for verification of
   this COSE object to a particular cryptographic algorithm.

9.2.1.  Signing of Registrar Voucher Request (RVR)

   A Registrar MUST include a COSE 'x5bag' structure in the RVR as
   explained in Section 8.1.  Below, an example Registrar Voucher
   Request (RVR) is shown that includes the 'x5bag' unprotected header
   parameter (32).  The bag contains two certificates in this case.

   18(                    / tag for COSE_Sign1                      /
    [
      h'A10126',          / protected COSE header encoding: {1: -7} /
                          /            which means { "alg": ES256 } /
      {                   / unprotected COSE header/
        32: [h'308202' ... '20AE', h'308201' ... '8CFF']    / x5bag /
      },
      h'A178' ... '7FED', / RVR payload wrapped in CBOR byte string /
      h'E1B7' ... '2925'  / RVR binary Sign1 signature              /
    ]
   )

   Besides storing the Registrar's own RVR-signing certificate chain
   (per Section 8.1), the Registrar MUST include in the same 'x5bag'
   structure all the certificates that the Pledge used to identify
   itself in the Pledge/Registrar DTLS handshake, including the End-
   Entity (IDevID) certificate and all CAs up to the root CA.  This
   serves two purposes:

   1.  A MASA that does not store the IDevIDs for all Pledges and their
       related sub-CAs is still able to reconstruct the certificate
       chain for a given Pledge and validate the Pledge's signature on
       the PVR based purely on the root CA of the Pledge's manufacturer
       that the MASA is storing.






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   2.  Diagnostic/debugging purposes: since the PVR's COSE signature
       does not store any certificates related to the signer, but only
       the signature itself, it can be useful for the MASA to log or
       inspect the Pledge's certificate chain in case the onboarding
       attempt fails.  Having the complete signing certificate chain at
       hand facilitates finding the root cause of the problem and helps
       in the communication with the customer.

   A 'kid' (key ID) field is OPTIONAL in the unprotected COSE header
   parameters map of a COSE object.  If present, it identifies the
   public key of the key pair that was used to sign the COSE message.
   The value of the key identifier 'kid' parameter may be in any format
   agreed between signer and verifier.  Usually a hash of the public key
   is used to identify the public key; but the choice of key identifier
   method is vendor-specific.

   By default, a Registrar does not include a 'kid' parameter in the RVR
   since the signing key is already identified by the signing
   certificates chain included in the COSE 'x5bag' structure.  A
   Registrar nevertheless MAY use a 'kid' parameter in its RVR to
   identify its signing key/identity.

   The method of generating such 'kid' value is vendor-specific and
   SHOULD be configurable in the Registrar to support commonly used
   methods.  In order to support future business cases and supply chain
   integrations, a Registrar using the 'kid' field MUST be configurable,
   on a per-manufacturer basis, to select a particular method for
   generating the 'kid' value such that it is compatible with the method
   that the manufacturer expects.  Note that the 'kid' field always has
   a CBOR byte string (bstr) format.

   In Appendix C.4 a further example of a signed RVR is shown.

9.2.2.  Signing of Pledge Voucher Request (PVR)

   Like in the RVR, a 'kid' (key ID) field is also OPTIONAL in the PVR.
   It can be used to identify the signing key/identity to the MASA.

   A Pledge by default SHOULD NOT use a 'kid' parameter in its PVR,
   because its signing key is already identified by the Pledge's unique
   serial number that is included in the PVR and (by the Registrar) in
   the RVR.  This achieves the smallest possible PVR data size while
   still enabling the MASA to fully verify the PVR.  Still, when
   required the Pledge MAY use a 'kid' parameter in its PVR to help the
   MASA identify the right public key to verify against.  This can occur
   for example if a Pledge has multiple IDevID identities.  The 'kid'
   parameter in this case may be an integer byte identifying one out of
   N identities present, or it may be a hash of the public key, or



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   anything else the Pledge vendor decides.  A Registrar normally SHOULD
   ignore a 'kid' parameter used in a received PVR, as this information
   is intended for the MASA.  In other words, there is no need for the
   Registrar to verify the contents of this field, but it may include it
   in an audit log.

   The example below shows a PVR with 'kid' present as an unprotected
   header parameter.

   18(                    / tag for COSE_Sign1                      /
    [
      h'A10126',          / protected COSE header encoding: {1: -7} /
                          /            which means { "alg": ES256 } /
      {
         4: h'59AB3E'     / COSE "kid" header parameter             /
      },
      h'A119' ... '3839', / PVR payload wrapped in CBOR byte string /
      h'5678' ... '7890'  / PVR binary Sign1 signature              /
    ]
   )

   The Pledge SHOULD NOT use the 'x5bag' or 'x5chain' COSE header
   parameters in the PVR.  A Registrar that processes a PVR with such a
   structure MUST ignore it, and MUST use only the TLS Client
   Certificate extension for authentication of the Pledge.

   A situation where the Pledge MAY use the 'x5bag' or 'x5chain'
   structure is for communication of certificate chains to the MASA.
   This would arise in some vendor- specific situations involving
   outsourcing of MASA functionality, or rekeying of the IDevID
   certification authority.

   In Appendix C.3 a further example of a signed PVR is shown.

9.2.3.  Signing of Voucher by MASA

   The MASA SHOULD NOT use a 'kid' parameter in the voucher response,
   because the MASA's signing key is already known to the Pledge.
   Still, where needed the MASA MAY use a 'kid' parameter in the voucher
   response to help the Pledge identify the right MASA public key to
   verify against.  This can occur for example if a Pledge has multiple
   IDevID identities.

   The MASA SHOULD NOT include an 'x5bag' or 'x5chain' attribute in the
   protected header of the voucher response, because normally a Pledge
   already stores the required public key for validation of the signed
   voucher.  The exception case is if the MASA knows that the Pledge
   doesn't pre-store the MASA's public key used for signing, and thus



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   the MASA needs to provide a certificate or certificate chain that
   will enable linking the signing identity to a pre-stored Trust Anchor
   (CA) in the Pledge.  This approach is not recommended, because
   including certificates in the protected 'x5bag' or 'x5chain' COSE
   header parameters will significantly increase the size of the voucher
   which impacts cBRSKI operation on constrained networks.

   For example, if the MASA signing key is based upon a PKI (see
   [I-D.richardson-anima-masa-considerations] Section 2.3), and the
   Pledge only pre-stores a manufacturer (root) CA identity in its Trust
   Store which is not the identity that signs the voucher, then a
   certificate chain needs to be included with the voucher in order for
   the Pledge to validate the MASA signing CA's signature by validating
   the chain up to the CA in its Trust Store.  In BRSKI CMS signed
   vouchers [RFC8995], the CMS structure has a place for such a
   certificate chain.  In cBRSKI COSE-signed vouchers, the 'x5bag'
   attribute [RFC9360] placed in the COSE protected header parameters is
   used to contain the needed certificates for the Pledge to form the
   chain.

   To signal the complete chain of the MASA's signing identity to the
   Registrar, the MASA MUST include the complete chain of signing
   certificates in an 'x5bag' attribute in the unprotected header of the
   voucher.  This allows the Registrar to optionally validate the
   voucher before forwarding it to the Pledge, or to validate it for
   logging purposes.  There is no voucher size impact of including this
   certificate chain in an unprotected 'x5bag' COSE header parameter for
   constrained networks, because the Registrar will remove this
   unprotected attribute prior to forwarding the voucher response to the
   Pledge, as defined in Section 6.6.

   Note that cBRSKI currently does not support signing the voucher with
   an RPK for which there is no corresponding certificate at all.  If
   the MASA wants to sign a voucher with an RPK that is not part of any
   PKIX hierarchy, it creates a single self-signed "placeholder" root CA
   certificate that uses the designated RPK as the public key.  This
   "placeholder" certificate is then included as the sole certificate in
   an unprotected 'x5bag' header parameter, as defined in the previous
   paragraph.

   Below, an example is shown of a COSE-signed voucher as created by
   MASA.  This example shows the common case where a protected 'x5bag'
   (32) attribute is not used, while an unprotected 'x5bag' (32)
   attribute is used to communicate the MASA's signature certificate
   chain to the Registrar.  The bag contains two certificates in this
   example.  One of these is the identity of the signer of the
   COSE_Sign1 object, whose signature is stored as the last CBOR array
   element in the below example.



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   18(                    / tag for COSE_Sign1                        /
    [
      h'A10126',          / protected COSE header encoding: {1: -7}   /
                          /            which means { "alg": ES256 }   /
      {                   / unprotected COSE header parameters        /
        32: [h'308202' ... '20AE', h'308201' ... '8CFF']      / x5bag /
      },
      h'A119' ... '3839', / voucher payload wrapped in CBOR byte str  /
      h'2A2C' ... '7FBF'  / voucher binary Sign1 signature by MASA    /
    ]
   )

   In Appendix C.5 a further example of a signed voucher is shown.

9.2.4.  Optional Validation of Voucher by Registrar

   For a Registrar, validation of the voucher and/or the signature of
   the voucher is optional, per Section 5.6 of [RFC8995].  However, if a
   Registrar does perform the validation of the signature chain,
   communicated in the 'x5bag' unprotected COSE header parameter (see
   Section 9.2.3)), it MUST validate that one of the below cases hold:

   1.  The signature chain is a single self-signed root CA certificate
       with a correct signature; and the public key in this certificate
       is also the public key that signed the voucher.  This represents
       the case where a voucher has been effectively signed with an RPK.

   2.  The signature chain consists of one or more certificates that can
       be chained to a known (preconfigured) trust root in the
       Registrar.

   The above validation elements are needed only for cases where
   (nonceless) vouchers are communicated over potentially unsecure
   channels to the Registrar.  Since the 'x5bag' header parameter is not
   protected by the voucher's COSE signature, it could have been
   modified in transit.

9.2.5.  Additional Information in the COSE Header

   The COSE header of the signed voucher can contain COSE header
   parameters with additional information, to be used by the Pledge.
   This information is additional to, and separate from, the voucher
   data defined by [RFC8366bis].

   An example of how this additional information can be used is adding a
   CBOR Web Token (CWT, [RFC8392]) claim in the COSE header as defined
   by [RFC9597], to encode the COSE signing time as an integer value in
   an 'iat' (Issued At) CWT claim.  This information in an integer



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   format may be useful for a Pledge that does not have date/time
   parsing functions, so it is unable to parse the date/time string
   value contained in the voucher.  Many other types of CWT claims can
   be included in a voucher in the same way, as needed for particular
   use cases.

   Such additional information can also be included in a COSE header of
   a voucher request by a Pledge, to be used by the MASA.

10.  Extensions to Discovery

   It is assumed that a Join Proxy (Section 6.2) seamlessly provides a
   relayed DTLS connection between the Pledge and the Registrar.  To use
   a Join Proxy, a Pledge needs to discover it.  For Pledge discovery of
   a Join Proxy, this section extends Section 4.1 of [RFC8995] for the
   cBRSKI case.

   In general, the Pledge may be one or more hops away from the
   Registrar, where one hop means the Registrar is a direct link-local
   neighbor of the Pledge.  The case of one hop away can be considered
   as a degenerate case, because a Join Proxy is not really needed then.

   The degenerate case would be unusual in constrained wireless network
   deployments, because a Registrar would typically not have a wireless
   network interface of the type used by constrained devices.  Rather,
   it would have a high-speed network interface.  Nevertheless, the
   situation where the Registrar is one hop away from the Pledge could
   occur in various cases like wired IoT networks or in wireless
   constrained networks where the Pledge is in radio range of a 6LoWPAN
   Border Router (6LBR) ([RFC6775])and the 6LBR happens to host a
   Registrar.

   In order to support the degenerate case, the Registrar SHOULD
   announce itself as if it were a Join Proxy -- though it would
   actually announce its real (stateful) Registrar CoAPS endpoint.  No
   actual Join Proxy functionality is then required on the Registrar.

   That way, a Pledge only needs to discover a Join Proxy, regardless of
   whether it is one or more than one hop away from a relevant
   Registrar.  It first discovers the link-local address and the UDP
   join-port of a Join Proxy.  The Pledge then follows the cBRSKI
   procedure of initiating a DTLS connection using the link-local
   address and join-port of the Join Proxy.








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   Once enrolled, a Pledge itself may function as a Join Proxy.  The
   decision whether or not to provide this functionality depends upon
   many factors and is out of scope for this document.  Such a decision
   might depend upon the amount of energy available to the device, the
   network bandwidth available, as well as CPU and memory availability.

   The process by which a Pledge discovers the Join Proxy, and how a
   Join Proxy discovers the location of the Registrar, are the subject
   of the remainder of this section.  Further details on both these
   topics are provided in [I-D.ietf-anima-constrained-join-proxy].

10.1.  Discovery Operations by a Pledge

   The Pledge must discover the address/port and optionally the protocol
   with which to communicate.  The present document only defines coaps
   (CoAP over DTLS) as the default protocol for cBRSKI, therefore
   protocol discovery is out of scope.

   For the discovery method, this section only defines unsecured CoAP
   discovery per Section 7 of [RFC7252] as the default method.  This
   uses CoRE Link Format [RFC6690] payloads.

   Section 11 briefly mentions other methods that apply to specific
   deployment types or technologies.  Details about these deployment-
   specific methods, or yet other methods, new payload formats, or more
   elaborate CoAP-based methods, may be defined in future documents such
   as [I-D.ietf-anima-brski-discovery].  The more elaborate methods for
   example may include discovering only Join Proxies that support a
   particular desired onboarding protocol, voucher format, or cBRSKI
   variant.

   Note that identifying the format of the voucher request and the
   voucher is currently not a required part of the Pledge's discovery
   operation.  It is assumed that all Registrars support all relevant
   voucher(-request) formats, while the Pledge only supports a single
   format.  A Pledge that makes a voucher request to a Registrar that
   does not support that format will receive a CoAP 4.06 Not Acceptable
   status code and the onboarding attempt will fail.

   Using CoAP discovery, a Pledge can discover a Join Proxy by sending a
   link-local multicast discovery message to the All CoAP Nodes address
   FF02::FD.  Zero, one, or multiple Join Proxies may respond.  The
   handling of multiple responses and absence of responses cases follow
   the guidelines of Section 4 of [RFC8995].  The discovery message is a
   CoAP GET request on the URI path /.well-known/core using a URI query
   "rt=brski.jp".  This resource type ('rt') is defined in Section 8.1
   of [I-D.ietf-anima-constrained-join-proxy].




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   Responding Join Proxies return a CoRE Link Format document with one
   or more links.  Each link indicates one CoAPS endpoint that offers
   cBRSKI Join Proxy functionality.  Formally, each link indicates a
   CoAP root resource (/) tagged with the "brski.jp" type, hosted on a
   CoAPS endpoint.

   In case a Pledge selects a particular Join Proxy for cBRSKI
   onboarding, it MUST use the link-local source address of the Join
   Proxy's discovery response as the destination IP address for its
   subsequent onboarding attempt.  This implies that the UTF-8 encoded
   link-local address literal that appears in the host subcomponent of
   each returned link is not used for determining the destination IP
   address of the onboarding attempt.

10.1.1.  Examples

   Below, a typical example is provided showing the Pledge's CoAP
   request and the Join Proxy's CoAP response.  The Join Proxy responds
   with a link-local source address, which is the same address as
   indicated in the URI-reference element ([RFC6690]) in the link in the
   discovery response payload.  The Join Proxy has a dedicated UDP port
   8485 open for DTLS connections of Pledges:

     REQ: GET coap://[ff02::fd]/.well-known/core?rt=brski.jp

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         <coaps://[fe80::c78:e3c4:58a0:a4ad]:8485>;rt=brski.jp

   Note that the Pledge would use the port number from the link in this
   response, but not the IPv6 literal in the host subcomponent
   ([fe80::c78:e3c4:58a0:a4ad]), as defined by the above discovery
   operation steps.

   The next example shows a Join Proxy that uses the default CoAPS port
   5684 for DTLS connections of Pledges.  In this case, the Join Proxy
   host is not using port 5684 for any other purposes, so it has the
   port available exclusively for accepting DTLS connections of Pledges.

     REQ: GET coap://[ff02::fd]/.well-known/core?rt=brski.jp

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         <coaps://[fe80::c78:e3c4:58a0:a4ad]>;rt=brski.jp





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   In the following example, two Join Proxies respond to the multicast
   query.  The Join Proxies each use a slightly different CoRE Link
   Format 'rt' value encoding.  While the first encoding is more
   compact, both encodings are allowed per [RFC6690].  The Pledge may
   now select one of the two Join Proxies for initiating its DTLS
   connection.

     REQ: GET coap://[ff02::fd]/.well-known/core?rt=brski*

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         <coaps://[fe80::c78:e3c4:58a0:a4ad]:8485>;rt=brski.jp

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         <coaps://[fe80::d359:3813:f382:3b23]:63245>;rt="brski.jp"

   In the final example, a single Join Proxy host responds with two
   distinct cBRSKI endpoints.  The Pledge may now select one of the two
   CoAP endpoints for initiating its DTLS connection.

     REQ: GET coap://[ff02::fd]/.well-known/core?rt=brski*

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         <coaps://[fe80::d359:3813:f382:3b23]:61616>;rt=brski.jp
         <coaps://[fe80::d359:3813:f382:3b23]:61617>;rt=brski.jp;
                                                     var="c509 v2"

   The first endpoint on port 61616 supports only the cBRSKI protocol as
   defined by this document.  The second endpoint, on port 61617,
   supports the same cBRSKI protocol as well as additional variations or
   extensions.  In this example, these variations/extensions are encoded
   using string values in a single attribute 'var'.  This information
   may be also encoded using other attributes defined by a future
   specification.

   A Pledge not aware of these variations can safely ignore these
   values, because the base cBRSKI protocol is supported by both
   endpoints, as indicated by the resource type ('rt').  If however a
   Pledge is aware of these variations, it can select the endpoint with
   the variation it prefers, in case multiple options are discovered.
   The use of attributes with a single base resource type allows future
   extensibility of cBRSKI, and enables the Join Proxies to support
   cBRSKI variants that are unknown to them.



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10.2.  Discovery Operations by a Join Proxy

   A Join Proxy needs to discover a Registrar, either at the moment it
   needs to relay data (of a Pledge) towards the Registrar, or prior to
   that moment.  For example, it may start Registrar discovery as soon
   as it is requested to be enabled in a Join Proxy role.  It may
   periodically redo this discovery, or periodically or on-demand check
   that the Registrar is still available in the network at the
   discovered IP address.

   As shown in the final example in Section 10.1.1, a Join Proxy can
   discover multiple Registrars in its network and present these options
   to the Pledge.  Each of these Registrars may support specific
   variations/extensions of cBRSKI - which may be defined in future
   documents.  It is up to the administrator of the network how many
   Registrars are enabled.

   Further details on CoAP discovery of the Registrar by a Join Proxy
   are provided in Section 5.1.1 of
   [I-D.ietf-anima-constrained-join-proxy].

11.  Deployment-specific Discovery Considerations

   This section details how discovery of a Join Proxy is done by the
   Pledge in specific deployment scenarios.  Future work such as
   [I-D.ietf-anima-brski-discovery] may define more details on discovery
   operations in the specific deployments listed here.

11.1.  6TiSCH Deployments

   In 6TiSCH networks, the Constrained Join Protocol (CoJP) is used as
   described in [RFC9031].  Such networks are expected to use EDHOC
   [RFC9528] for key management, which is out of scope of this document.
   The IEEE 802.15.4 Enhanced Beacon has been extended in [RFC9032] to
   allow for discovery of a 6TiSCH-compliant Join Proxy.

11.2.  IP networks using GRASP

   In IP networks that support GRASP [RFC8990], a Pledge can discover a
   Join Proxy by listening for GRASP messages.  GRASP supports mesh
   networks, and can also be used over unencrypted Wi-Fi.

   Details of GRASP discovery of constrained Join Proxies are out of
   scope of this document and may be defined in future work.







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11.3.  IP networks using mDNS

   [RFC8995] defines a mechanism for the Pledge to discover a Join Proxy
   by sending mDNS [RFC6762] queries.  This mechanism can be used on any
   IP network which does not have another recommended mechanism.  It can
   be used over unencrypted Wi-Fi.  This mechanism does support link-
   local Join Proxy discovery in mesh networks.  However, it does not
   support Registrar discovery by Join Proxies in mesh networks, because
   the Registrar is typically not reachable by link-local communication
   in that case.  For this, another mechanism is needed, which is out of
   scope of this document and may be defined in future work.

   A Pledge uses an mDNS PTR query for the name "_brski-
   proxy._udp.local." to discover link-local constrained Join Proxies.
   The label "_udp" here indicates a query for cBRSKI constrained Join
   Proxies, as opposed to "_tcp" defined in [RFC8995] which is for
   discovering BRSKI Proxies.

11.4.  Thread Networks using Mesh Link Establishment (MLE)

   Thread [Thread] is a wireless mesh network protocol based on 6LoWPAN
   [RFC6282] and other IETF protocols.  In Thread, a new device
   discovers potential Thread networks and Thread nodes to join by using
   the Mesh Link Establishment (MLE)
   [I-D.ietf-6lo-mesh-link-establishment] protocol.  MLE uses the UDP
   port number 19788.

   The new device sends discovery requests on different IEEE 802.15.4
   radio channels, to which Thread nodes (if any present) respond with a
   discovery response containing information about their respective
   network.  The MLE discovery response message contains UDP port
   information to signal the new device which UDP port to use for its
   DTLS connection to the Join Proxy function.  The link-local IPv6
   source address of the MLE response message indicates the address of
   the Join Proxy.

   Once a suitable Thread node is selected as its Join Proxy, the new
   device initiates a DTLS transport-layer secured connection to the
   network's commissioning application, over a link-local single radio
   hop to the selected Join Proxy.  This link is not yet secured at the
   radio/MAC link layer: link-layer security will be set up once the new
   device is approved by the commissioning application to join the
   Thread network, and it gets provisioned with network access
   credentials.

   A Thread node that is capable to act as a Join Proxy will only enable
   this role if the network-wide configuration data indicates that new
   device commissioning is allowed.



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12.  Design and Implementation Considerations

12.1.  Voucher Format and Encoding

   The design considerations for vouchers from Section 10 of
   [RFC8366bis] apply.  Specifically for CBOR encoding of voucher data,
   one key difference with JSON encoding is that the names of the leaves
   in the YANG definition do not affect the size of the resulting CBOR,
   if the SID ([RFC9254], [RFC9595]) translation process is used that
   assigns integers to the names.

   To obtain the lowest code size and RAM use on the Pledge, it is
   recommended that a Pledge is designed to only process/generate these
   SID integers and not the lengthy strings.  The MASA in that case is
   required to generate the voucher data for that Pledge using only SID
   integers.  Yet, this MASA MUST still support both SID integers and
   strings, to be able to process attribute (string) names in the RVR
   which the Registrar may use.

12.2.  CoAP Usage

   A successful POST request to the Registrar's telemetry resources
   (/vs, /es) returns a 2.04 Changed response with empty payload.

   A CoAP client sending a request should be aware that the server, even
   in case of an empty payload, may use either a piggybacked CoAP
   response (for example ACK with code 2.04) but may also respond with a
   separate CoAP response.  This is first an ACK message with code 0.0
   that acknowledges the reception of the request.  It is followed by a
   CON message with a code 2.04 response in a separate CoAP message.
   See [RFC7252] for details.

12.3.  Use of cBRSKI with HTTPS

   This specification contains two major extensions to [RFC8995]: a
   constrained voucher format (COSE), and a constrained transfer
   protocol (CoAP).

   On constrained networks with constrained devices, it make senses to
   use both together.  However, this document does not mandate that this
   be the only way.

   A given constrained device design and software may be re-used for
   multiple device models, such as a model having only an IEEE 802.15.4
   radio, or a model having only an IEEE 802.11 (Wi-Fi) radio, or a
   model having both these radios.  A manufacturer of such device models
   may wish to have code only for the use of the constrained voucher
   format (COSE), and use it on all supported radios including the IEEE



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   802.11 radio.  For this radio, the software stack to support HTTP/TLS
   may be already integrated into the radio module hence it is
   attractive for the manufacturer to reuse this.  This type of approach
   is supported by this document.  In the case that HTTPS is used, the
   regular long [RFC8995] resource names are used, together with the new
   application/voucher+cose media type described in this document.  For
   status telemetry requests, the format and requirements defined in
   Section 6.3.1 remain unchanged.

   Other combinations are possible, but they are not enumerated here.
   New work such as [I-D.ietf-anima-jws-voucher] provides new formats
   that may be usable over a number of different transports.  In
   general, sending larger payloads over constrained networks makes less
   sense, while sending smaller payloads over unconstrained networks is
   perfectly acceptable.

   The Pledge will in most cases support a single voucher format, which
   it uses without negotiation i.e. without discovery of formats
   supported.  The Registrar, being unconstrained, is expected to
   support all voucher formats.  There will be cases where a Registrar
   does not support a new format that a new Pledge uses, and this is an
   unfortunate situation that will result in lack of interoperation.

   The responsibility for supporting new formats is on the Registrar.

13.  Raw Public Key Variant

13.1.  Introduction and Scope

   This section introduces a cBRSKI variant to further reduce the data
   volume and complexity of the cBRSKI onboarding.  The use of a raw
   public key (RPK) in the pinning process can significantly reduce the
   number of bytes sent over the wire and the number of round trips, and
   reduce the code footprint in a Pledge.  But it comes with a few
   significant operational limitations.

   One simplification that comes with RPK use is that a Pledge can avoid
   doing PKIX operations, such as certificate chain validation.

13.2.  DTLS Connection and Registrar Trust Anchor

   When the Pledge first connects to the Registrar, the connection to
   the Registrar is provisional, as explained in Section 5.6.2 of
   [RFC8995].  The Registrar normally provides its public key in a
   TLSServerCertificate, and the Pledge uses that to validate that
   integrity of the DTLS connection, but it does not validate the
   identity of the provided certificate.




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   As the TLSServerCertificate object is never verified directly by the
   Pledge, sending it can be considered superfluous.  So instead of
   using a (TLSServer)Certificate of type X509 (see Section 4.4.2 of
   [RFC8446]), a RawPublicKey object (as defined by [RFC7250]) is used.

   A Registrar operating in a mixed environment can determine whether to
   send a PKIX certificate chain or a Raw Public Key to the Pledge: this
   is signaled by the Pledge.  In the case the Pledge needs an RPK, it
   includes the server_certificate_type of RawPublicKey.  This is shown
   in Section 5 of [RFC7250].

   The Pledge MUST send a client_certificate_type of X509 (not an RPK),
   so that the Registrar can properly identify the Pledge and distill
   the MASA URI information from its IDevID certificate.

13.3.  The Pledge Voucher Request

   The Pledge puts the Registrar's public key into the 'proximity-
   registrar-pubk' attribute of the Pledge Voucher Request (PVR).  The
   'proximity-registrar-pubk-sha256' can alternatively be used for
   efficiency, if the 32-bytes of a SHA256 hash turns out to be smaller
   than a typical ECDSA key.

   As the format of the 'proximity-registrar-pubk' attribute is
   identical to the TLS RawPublicKey data object, no manipulation at all
   is needed to insert this attribute into the PVR.  This approach
   reduces the size of the PVR significantly, compared to including the
   full certificate.

13.4.  The Voucher Response

   A returned voucher will have a 'pinned-domain-pubk' attribute with
   the identical key as was found in the 'proximity-registrar-pubk'
   attribute above, as well as being identical to the Registrar's RPK in
   the currently active DTLS connection.  (Or alternatively the MASA may
   include the 'pinned-domain-pubk-sha256' attribute if it knows the
   Pledge supports this attribute.)

   Validation of this key by the Pledge is what takes the DTLS
   connection out of the provisional state; see Section 5.6.2 of
   [RFC8995] for more details.

   The received voucher needs to be validated by the Pledge.  The Pledge
   needs to have a public key to validate the signature from the MASA on
   the voucher.






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   The MASA's public key counterpart of the (private) MASA signing key
   MUST be already installed in the Pledge at manufacturing time.
   Otherwise, the Pledge cannot validate the voucher's signature.

13.5.  The Enrollment Phase

   A Pledge that does not support PKIX operations cannot use EST to
   enroll; it has to use another method for enrollment without
   certificates and the Registrar has to support this method also.  For
   example, an enrollment process that records an RPK owned by the
   Pledge as a legitimate entity that is part of the domain.

   It is possible that the Pledge will not enroll after obtaining a
   valid voucher, but instead will do only a network join operation (see
   for example [RFC9031]).  How the Pledge discovers this method and
   details of such enrollment methods are out of scope of this document.

14.  Security Considerations

14.1.  Duplicate Serial Numbers

   If a manufacturer sold products with duplicated serial numbers, that
   use the same MASA CA as their root of trust, a customer of one of
   these products can potentially perform an attack where it uses a
   voucher created for product 1 to onboard product 2.  This attack only
   works for nonceless vouchers.

   Note that such a situation could only arise due to manufacturer mis-
   management or oversight.

   For example, imagine the same manufacturer makes light bulbs as well
   as gas centrifuges, and said manufacturer does not uniquely allocate
   product serial numbers.  The attacker has obtained a light bulb which
   happens to have the same serial number as an operational gas
   centrifuge which it wishes to obtain access to.  The attacker
   performs a normal BRSKI onboarding for the light bulb, but then uses
   the resulting nonceless voucher to onboard the gas centrifuge.  The
   attack requires that the gas centrifuge be returned to a state where
   it is willing to perform a new onboarding operation.  For example, a
   factory reset.

   This attack is normally prevented by the mechanisms of using
   different trust root CAs for different product lines, and/or using
   unique serial numbers within a single MASA CA scope.







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   Section 6.2 of [RFC8366bis] discusses cases of duplicated serial
   numbers in products across different CAs and the role of the 'idevid-
   issuer' attribute in the RVR and in the voucher to disambiguate these
   products.

14.2.  IDevID Security in the Pledge

   The security of this protocol depends upon the Pledge identifying
   itself to the Registrar using its manufacturer installed certificate:
   the IDevID certificate.  Associated with this certificate is the
   IDevID private key, known only to the Pledge.  Disclosure of this
   private key to an attacker would permit the attacker to impersonate
   the Pledge towards the Registrar, probably gaining access credentials
   to that Registrar's network.

   If the IDevID private key disclosure is known to the manufacturer,
   there is little recourse other than recall of the relevant part
   numbers.  The process for communicating this recall would be within
   the BRSKI-MASA protocol.  Neither this specification nor [RFC8995]
   provides for consultation of a Certification Revocation List (CRL) or
   Open Certificate Status Protocol (OCSP) by a Registrar when
   evaluating an IDevID certificate.  However, the BRSKI-MASA protocol
   submits the IDevID from the Registrar to the manufacturer's MASA and
   a manufacturer would have an opportunity to decline to issue a
   voucher for a device which they believe has become compromised.

   It may be difficult for a manufacturer to determine when an IDevID
   private key has been disclosed.  Two situations present themselves:
   in the first situation a compromised private key might be reused in a
   counterfeit device, which is sold to another customer.  This would
   present itself as an onboarding of the same device in two different
   networks.  The manufacturer may become suspicious seeing two voucher
   requests for the same device from different Registrars.  Such
   activity could be indistinguishable from a device which has been
   resold from one operator to another, or re-deployed by an operator
   from one location to another.















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   In the second situation, an attacker having compromised the IDevID
   private key of a device might then install malware into the same
   device and attempt to return it to service.  The device, now blank,
   would go through a second onboarding process with the original
   Registrar.  Such a Registrar could notice that the device has been
   "factory reset" and alert the operator to this situation.  One remedy
   against the presence of malware is through the use of Remote
   Attestation such as described in [RFC9334].  Future work will need to
   specify a background-check Attestation flow as part of the voucher
   request/response process.  Attestation may still require access to a
   private key (e.g. IDevID private key) in order to sign Evidence, so a
   primary goal should be to keep any private key safe within the
   Pledge.

   In larger, more expensive, systems there is budget (power, space, and
   bill of materials) to include more specific defenses for a private
   key.  For instance, this includes putting the IDevID private key in a
   Trusted Platform Module (TPM), or use of Trusted Execution
   Environments (TEE) for access to the key.  On smaller IoT devices,
   the cost and power budget for an extra part is often prohibitive.

   It is becoming more and more common for CPUs to have an internal set
   of one-time fuses that can be programmed (often they are "burnt" by a
   laser) at the factory.  This section of memory is only accessible in
   some privileged CPU state.  The use of this kind of CPU is
   appropriate as it provides significant resistance against key
   disclosure even when the device can be disassembled by an attacker.

   In a number of industry verticals, there is increasing concern about
   counterfeit parts.  These may be look-alike parts created in a
   different factory, or parts which are created in the same factory
   during an illegal night-shift, but which are not subject to the
   appropriate level of quality control.  The use of a manufacturer-
   signed IDevID certificate provides for discovery of the pedigree of
   each part, and this often justifies the cost of the security measures
   associated with storing the private key.

14.3.  Security of the BRSKI-MASA Protocol

   Section 7.1 explains why no CoAPS version of the BRSKI-MASA protocol
   is specified.  The connection from the Registrar to the MASA
   continues to be HTTPS as in [RFC8995].

   This choice enables the BRSKI-MASA protocol, which operates over the
   open Internet, to be secured using standard solutions that are
   commonly used for HTTPS over the Internet.  The use of UDP protocols
   across the Internet is sometimes fraught with security challenges.
   Denial-of-service attacks against UDP based protocols are trivial as



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   there is no three-way handshake as done for TCP.  The three-way
   handshake of TCP guarantees that the node sending the connection
   request is reachable using the origin IP address.  While DTLS
   contains an option to do a stateless challenge -- a process actually
   stronger than that done by TCP -- it is not yet common for this
   mechanism to be available in hardware at multigigabit speeds.

   Also, in many enterprise networks outgoing UDP connections can be
   treated as suspicious, which could effectively block CoAP connections
   for some firewall configurations.  Reducing the complexity of MASA
   (i.e. less protocols supported) also reduces its potential attack
   surface, which is relevant since the MASA is 24/7 exposed on the
   Internet and accepting (untrusted) incoming connections.

14.4.  Registrar Certificate May Be Self-signed

   The provisional (D)TLS connection formed by the Pledge with the
   Registrar does not authenticate the Registrar's identity.  This
   Registrar's identity is validated by the [RFC8366bis] voucher that is
   issued by the MASA, signed with a trust anchor that was built-in to
   the Pledge.

   The Registrar may therefore use any certificate, including a self-
   signed one.  The only restrictions on the certificate is that it MUST
   have EKU bits set as detailed in Section 6.1.5 and Section 7.4.

14.5.  Use of RPK Alternatives to 'proximity-registrar-cert'

   In Section 9 of [RFC8366bis] two compact alternative attributes for
   'proximity-registrar-cert' are defined that include an RPK:
   'proximity-registrar-pubk' and 'proximity-registrar-pubk-sha256'.
   The Pledge can use these attributes in its PVR to identify the
   Registrar based on its public key only.  Since the full certificate
   of the proximate Registrar is not included, use of these attributes
   by a Pledge implies that a Registrar could insert another certificate
   with the same public key identity into the RVR.  For example, an
   older or a newer version of its certificate.  The MASA will not be
   able to detect such act by the Registrar.  But since any certificate
   the Registrar could insert in this way still encodes its own identity
   the additional risk of using the RPK alternatives is negligible.

   When a Registrar sees a PVR that uses one of 'proximity-registrar-
   pubk' or 'proximity-registrar-pubk-sha256' attributes, this implies
   the Registrar must include the certificate identified by these
   attributes into its RVR.  Otherwise, the MASA is unable to verify
   proximity.  This requirement is already implied by the "MUST"
   requirement in Section 8.1.




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15.  IANA Considerations

15.1.  Resource Type Link Target Attribute Values Registry

   Additions to the "Resource Type (rt=) Link Target Attribute Values"
   IANA registry, within the "CoRE Parameters" registry group are
   specified below.

   Reference: [This RFC]

        +==========+==============================================+
        | Value    | Description                                  |
        +==========+==============================================+
        | brski    | Base resource of all Bootstrapping Remote    |
        |          | Secure Key Infrastructure (cBRSKI) resources |
        +----------+----------------------------------------------+
        | brski.rv | cBRSKI request voucher resource              |
        +----------+----------------------------------------------+
        | brski.vs | cBRSKI voucher status telemetry resource     |
        +----------+----------------------------------------------+
        | brski.es | cBRSKI enrollment status telemetry resource  |
        +----------+----------------------------------------------+
        | ace.est  | Base resource of all Enrollment over Secure  |
        |          | Transport CoAPS (EST-coaps) resources        |
        +----------+----------------------------------------------+

             Table 2: Resource Type (rt) link target attribute
                      values for cBRSKI and EST- coaps

   Note that the resource type "brski" identifies a base resource in a
   resource hierarchy on a CoAP server, where its sub-resources each
   have one of the resource types "brski.*" as defined by this
   specification.  Similarly, the resource type "ace.est" identifies a
   base resource in a resource hierarchy, where its sub-resources each
   have one of the resource types "ace.est.*" as defined by [RFC9148].

15.2.  Media Types Registry

   This section registers the media type application/voucher+cose in the
   "Media Types" IANA registry.  This media type is used to indicate
   that the content is a CBOR voucher or voucher request signed with a
   COSE_Sign1 structure [RFC9052] as defined in this document.

15.2.1.  application/voucher+cose







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   Type name:  application
   Subtype name:  voucher+cose
   Required parameters:  N/A
   Optional parameters:  N/A
   Encoding considerations:  binary (CBOR)
   Security considerations:  Section 14 of [This RFC], and Section 12
     of [RFC 9052] for the COSE_Sign1 structured that is used.
   Interoperability considerations:  Section 15.2.2 of [This RFC]
   Published specification:  [This RFC]
   Applications that use this media type:  cBRSKI/ANIMA, 6TiSCH, and
     other zero-touch onboarding systems
   Fragment identifier considerations: N/A
   Additional information:
     Deprecated alias names for this type: N/A
     Magic number(s):  N/A
     File extension(s):  .vch
     Macintosh file type code(s):  N/A
   Person & email address to contact for further information:  IETF
     ANIMA Working Group (anima@ietf.org) or IETF Operations and
     Management Area Working Group (opsawg@ietf.org)
   Intended usage:  COMMON
   Restrictions on usage:  N/A
   Author:  ANIMA WG
   Change controller:  IETF
   Provisional registration? (standards tree only):  NO

15.2.2.  Interoperability Considerations for application/voucher+cose

   The media type defined here does not have any parameter to indicate
   whether names are used, or SID integers are used, or both can be
   mixed within a voucher data item.  In absence of any specific further
   knowledge about this, a mixed use of SID integers and names MUST be
   assumed, which is equivalent to the application/yang-data+cbor media
   type ([RFC9254]) without the optional 'id' parameter.

   Furthermore,

   *  a Registrar assumes that mixed SIDs/names MAY be present in a
      received PVR or voucher;

   *  a MASA assumes that mixed SIDs/names MAY be present in a received
      RVR;

   *  a Pledge assumes, depending on its implementation, that SIDs are
      present only, or names are present only, or mixed SIDs/names are
      present in a received voucher.





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   Because the MASA and Pledge are under control (either directly or by
   contract) of the same manufacturer, they can be co-developed
   regarding the type of identifiers produced and identifiers consumed
   in order to guarantee interoperability.

15.3.  CoAP Content-Formats Registry

   IANA has allocated ID 836 from the "CoAP Content-Formats" registry as
   shown below.

        +==========================+==========+=====+============+
        | Media type               | Encoding | ID  | Reference  |
        +==========================+==========+=====+============+
        | application/voucher+cose | -        | 836 | [This RFC] |
        +--------------------------+----------+-----+------------+

           Table 3: Additions to the IANA CoAP Content-Formats
                                 Registry

15.4.  Update to BRSKI Well-Known URIs Registry

   This section updates the "BRSKI Well-Known URIs" IANA registry of the
   Bootstrapping Remote Secure Key Infrastructures (BRSKI) Parameters
   Registry group, by adding a new column "Short Path Segment",
   clarifying existing "Description" values, and renaming the column
   "URI" to "Path Segment".

   The new "Short Path Segment" entry denotes a shorter alternative to
   Path Segment for the resource that can be used by a client in a CoAP
   request on a well-known BRSKI resource.  A value "N/A" can be
   registered to denote that there is no short path segment defined.

   The contents of the registry with these changes applied are as
   follows:

















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   +=================+============+=======================+============+
   | Path Segment    | Short      | Description           | Reference  |
   |                 | Path       |                       |            |
   |                 | Segment    |                       |            |
   +=================+============+=======================+============+
   | requestvoucher  | rv         | Request voucher:      | [RFC8995], |
   |                 |            | Pledge to Registrar,  | [This RFC] |
   |                 |            | and Registrar to MASA |            |
   +-----------------+------------+-----------------------+------------+
   | voucher_status  | vs         | Voucher status        | [RFC8995], |
   |                 |            | telemetry: Pledge to  | [This RFC] |
   |                 |            | Registrar             |            |
   +-----------------+------------+-----------------------+------------+
   | requestauditlog | N/A        | Request audit log:    | [RFC8995]  |
   |                 |            | Registrar to MASA     |            |
   +-----------------+------------+-----------------------+------------+
   | enrollstatus    | es         | Enrollment status     | [RFC8995], |
   |                 |            | telemetry: Pledge to  | [This RFC] |
   |                 |            | Registrar             |            |
   +-----------------+------------+-----------------------+------------+

         Table 4: Update of the IANA BRSKI Well-Known URIs Registry

15.5.  Structured Syntax Suffixes Registry

   This section registers the "+cose" suffix in the "Structured Syntax
   Suffixes" IANA Registry based on the [RFC6838] procedure.
























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   Name:       CBOR Object Signing and Encryption (COSE) object
   +suffix:    +cose
   References: the application/cose media type [RFC9052]
   Encoding considerations: binary (CBOR)
   Interoperability considerations:
     the application/cose media type has an optional parameter
     "cose-type". Any new media type that uses the +cose suffix
     and allows use of this parameter MUST specify this
     explicitly, per Section 4.3 of [RFC6838]. If the parameter
     "cose-type" is allowed, its usage MUST be identical to the
     usage defined for the application/cose media type in
     Section 2 of [RFC9052].
     A COSE processor handling a media type foo+cose and which
     does not know the specific type "foo" SHOULD use the
     cose-type COSE tag, if present, or cose-type parameter, if
     present, to determine the specific COSE object type during
     processing. If the specific type cannot be determined,
     it MUST assume only the generic COSE object structure and
     it MUST NOT perform security-critical operations using the
     COSE object.
   Fragment identifier considerations: N/A
   Security considerations: see [RFC9052]
   Contact:
     IETF COSE Working Group (cose@ietf.org) or
     IESG (iesg@ietf.org)
   Author/Change controller:
     IETF ANIMA Working Group (anima@ietf.org).
     IESG has change control over this registration.

16.  References

16.1.  Normative References

   [I-D.ietf-anima-constrained-join-proxy]
              Dijk, E., Richardson, M., Van der Stok, P., and P.
              Kampanakis, "Join Proxy for Bootstrapping of Constrained
              Network Elements", Work in Progress, Internet-Draft,
              draft-ietf-anima-constrained-join-proxy-18, 19 October
              2025, <https://datatracker.ietf.org/doc/html/draft-ietf-
              anima-constrained-join-proxy-18>.

   [I-D.ietf-core-uri-path-abbrev]
              Amsüss, C. and M. Richardson, "URI-Path abbreviation in
              CoAP", Work in Progress, Internet-Draft, draft-ietf-core-
              uri-path-abbrev-02, 20 October 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-core-
              uri-path-abbrev-02>.




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   [ieee802-1AR]
              "IEEE 802.1AR Secure Device Identity", IEEE Standards
              Association, 2018,
              <https://standards.ieee.org/ieee/802.1AR/6995/>.

   [RFC2119]  Bradner, S., "Key words for use in RFCs to Indicate
              Requirement Levels", BCP 14, RFC 2119,
              DOI 10.17487/RFC2119, March 1997,
              <https://www.rfc-editor.org/rfc/rfc2119>.

   [RFC4193]  Hinden, R. and B. Haberman, "Unique Local IPv6 Unicast
              Addresses", RFC 4193, DOI 10.17487/RFC4193, October 2005,
              <https://www.rfc-editor.org/rfc/rfc4193>.

   [RFC4210]  Adams, C., Farrell, S., Kause, T., and T. Mononen,
              "Internet X.509 Public Key Infrastructure Certificate
              Management Protocol (CMP)", RFC 4210,
              DOI 10.17487/RFC4210, September 2005,
              <https://www.rfc-editor.org/rfc/rfc4210>.

   [RFC5280]  Cooper, D., Santesson, S., Farrell, S., Boeyen, S.,
              Housley, R., and W. Polk, "Internet X.509 Public Key
              Infrastructure Certificate and Certificate Revocation List
              (CRL) Profile", RFC 5280, DOI 10.17487/RFC5280, May 2008,
              <https://www.rfc-editor.org/rfc/rfc5280>.

   [RFC6066]  Eastlake 3rd, D., "Transport Layer Security (TLS)
              Extensions: Extension Definitions", RFC 6066,
              DOI 10.17487/RFC6066, January 2011,
              <https://www.rfc-editor.org/rfc/rfc6066>.

   [RFC6347]  Rescorla, E. and N. Modadugu, "Datagram Transport Layer
              Security Version 1.2", RFC 6347, DOI 10.17487/RFC6347,
              January 2012, <https://www.rfc-editor.org/rfc/rfc6347>.

   [RFC6690]  Shelby, Z., "Constrained RESTful Environments (CoRE) Link
              Format", RFC 6690, DOI 10.17487/RFC6690, August 2012,
              <https://www.rfc-editor.org/rfc/rfc6690>.

   [RFC6762]  Cheshire, S. and M. Krochmal, "Multicast DNS", RFC 6762,
              DOI 10.17487/RFC6762, February 2013,
              <https://www.rfc-editor.org/rfc/rfc6762>.

   [RFC7030]  Pritikin, M., Ed., Yee, P., Ed., and D. Harkins, Ed.,
              "Enrollment over Secure Transport", RFC 7030,
              DOI 10.17487/RFC7030, October 2013,
              <https://www.rfc-editor.org/rfc/rfc7030>.




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   [RFC7250]  Wouters, P., Ed., Tschofenig, H., Ed., Gilmore, J.,
              Weiler, S., and T. Kivinen, "Using Raw Public Keys in
              Transport Layer Security (TLS) and Datagram Transport
              Layer Security (DTLS)", RFC 7250, DOI 10.17487/RFC7250,
              June 2014, <https://www.rfc-editor.org/rfc/rfc7250>.

   [RFC7252]  Shelby, Z., Hartke, K., and C. Bormann, "The Constrained
              Application Protocol (CoAP)", RFC 7252,
              DOI 10.17487/RFC7252, June 2014,
              <https://www.rfc-editor.org/rfc/rfc7252>.

   [RFC7959]  Bormann, C. and Z. Shelby, Ed., "Block-Wise Transfers in
              the Constrained Application Protocol (CoAP)", RFC 7959,
              DOI 10.17487/RFC7959, August 2016,
              <https://www.rfc-editor.org/rfc/rfc7959>.

   [RFC8174]  Leiba, B., "Ambiguity of Uppercase vs Lowercase in RFC
              2119 Key Words", BCP 14, RFC 8174, DOI 10.17487/RFC8174,
              May 2017, <https://www.rfc-editor.org/rfc/rfc8174>.

   [RFC8366bis]
              Watsen, K., Richardson, M., Dijk, E., Eckert, T. T., and
              Q. Ma, "A Voucher Artifact for Bootstrapping Protocols",
              Work in Progress, Internet-Draft, draft-ietf-anima-
              rfc8366bis-27, 26 February 2026,
              <https://datatracker.ietf.org/doc/html/draft-ietf-anima-
              rfc8366bis-27>.

   [RFC8446]  Rescorla, E., "The Transport Layer Security (TLS) Protocol
              Version 1.3", RFC 8446, DOI 10.17487/RFC8446, August 2018,
              <https://www.rfc-editor.org/rfc/rfc8446>.

   [RFC8449]  Thomson, M., "Record Size Limit Extension for TLS",
              RFC 8449, DOI 10.17487/RFC8449, August 2018,
              <https://www.rfc-editor.org/rfc/rfc8449>.

   [RFC8610]  Birkholz, H., Vigano, C., and C. Bormann, "Concise Data
              Definition Language (CDDL): A Notational Convention to
              Express Concise Binary Object Representation (CBOR) and
              JSON Data Structures", RFC 8610, DOI 10.17487/RFC8610,
              June 2019, <https://www.rfc-editor.org/rfc/rfc8610>.

   [RFC8710]  Fossati, T., Hartke, K., and C. Bormann, "Multipart
              Content-Format for the Constrained Application Protocol
              (CoAP)", RFC 8710, DOI 10.17487/RFC8710, February 2020,
              <https://www.rfc-editor.org/rfc/rfc8710>.





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   [RFC8949]  Bormann, C. and P. Hoffman, "Concise Binary Object
              Representation (CBOR)", STD 94, RFC 8949,
              DOI 10.17487/RFC8949, December 2020,
              <https://www.rfc-editor.org/rfc/rfc8949>.

   [RFC8995]  Pritikin, M., Richardson, M., Eckert, T., Behringer, M.,
              and K. Watsen, "Bootstrapping Remote Secure Key
              Infrastructure (BRSKI)", RFC 8995, DOI 10.17487/RFC8995,
              May 2021, <https://www.rfc-editor.org/rfc/rfc8995>.

   [RFC9031]  Vučinić, M., Ed., Simon, J., Pister, K., and M.
              Richardson, "Constrained Join Protocol (CoJP) for 6TiSCH",
              RFC 9031, DOI 10.17487/RFC9031, May 2021,
              <https://www.rfc-editor.org/rfc/rfc9031>.

   [RFC9032]  Dujovne, D., Ed. and M. Richardson, "Encapsulation of
              6TiSCH Join and Enrollment Information Elements",
              RFC 9032, DOI 10.17487/RFC9032, May 2021,
              <https://www.rfc-editor.org/rfc/rfc9032>.

   [RFC9052]  Schaad, J., "CBOR Object Signing and Encryption (COSE):
              Structures and Process", STD 96, RFC 9052,
              DOI 10.17487/RFC9052, August 2022,
              <https://www.rfc-editor.org/rfc/rfc9052>.

   [RFC9053]  Schaad, J., "CBOR Object Signing and Encryption (COSE):
              Initial Algorithms", RFC 9053, DOI 10.17487/RFC9053,
              August 2022, <https://www.rfc-editor.org/rfc/rfc9053>.

   [RFC9147]  Rescorla, E., Tschofenig, H., and N. Modadugu, "The
              Datagram Transport Layer Security (DTLS) Protocol Version
              1.3", RFC 9147, DOI 10.17487/RFC9147, April 2022,
              <https://www.rfc-editor.org/rfc/rfc9147>.

   [RFC9148]  van der Stok, P., Kampanakis, P., Richardson, M., and S.
              Raza, "EST-coaps: Enrollment over Secure Transport with
              the Secure Constrained Application Protocol", RFC 9148,
              DOI 10.17487/RFC9148, April 2022,
              <https://www.rfc-editor.org/rfc/rfc9148>.

   [RFC9254]  Veillette, M., Ed., Petrov, I., Ed., Pelov, A., Bormann,
              C., and M. Richardson, "Encoding of Data Modeled with YANG
              in the Concise Binary Object Representation (CBOR)",
              RFC 9254, DOI 10.17487/RFC9254, July 2022,
              <https://www.rfc-editor.org/rfc/rfc9254>.






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   [RFC9360]  Schaad, J., "CBOR Object Signing and Encryption (COSE):
              Header Parameters for Carrying and Referencing X.509
              Certificates", RFC 9360, DOI 10.17487/RFC9360, February
              2023, <https://www.rfc-editor.org/rfc/rfc9360>.

   [RFC9525]  Saint-Andre, P. and R. Salz, "Service Identity in TLS",
              RFC 9525, DOI 10.17487/RFC9525, November 2023,
              <https://www.rfc-editor.org/rfc/rfc9525>.

16.2.  Informative References

   [COSE-registry]
              IANA, "CBOR Object Signing and Encryption (COSE) registry
              group", 11 January 2017,
              <https://www.iana.org/assignments/cose/cose.xhtml>.

   [I-D.ietf-6lo-mesh-link-establishment]
              Kelsey, R., "Mesh Link Establishment", Work in Progress,
              Internet-Draft, draft-ietf-6lo-mesh-link-establishment-00,
              1 December 2015, <https://datatracker.ietf.org/doc/html/
              draft-ietf-6lo-mesh-link-establishment-00>.

   [I-D.ietf-anima-brski-discovery]
              Eckert, T. T. and E. Dijk, "BRSKI discovery and
              variations", Work in Progress, Internet-Draft, draft-ietf-
              anima-brski-discovery-09, 20 October 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-anima-
              brski-discovery-09>.

   [I-D.ietf-anima-jws-voucher]
              Werner, T. and M. Richardson, "JWS signed Voucher
              Artifacts for Bootstrapping Protocols", Work in Progress,
              Internet-Draft, draft-ietf-anima-jws-voucher-16, 15
              January 2025, <https://datatracker.ietf.org/doc/html/
              draft-ietf-anima-jws-voucher-16>.

   [I-D.ietf-cbor-edn-literals]
              Bormann, C., "CBOR Extended Diagnostic Notation (EDN)",
              Work in Progress, Internet-Draft, draft-ietf-cbor-edn-
              literals-19, 16 October 2025,
              <https://datatracker.ietf.org/doc/html/draft-ietf-cbor-
              edn-literals-19>.









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   [I-D.ietf-cose-cbor-encoded-cert]
              Mattsson, J. P., Selander, G., Raza, S., Höglund, J., and
              M. Furuhed, "CBOR Encoded X.509 Certificates (C509
              Certificates)", Work in Progress, Internet-Draft, draft-
              ietf-cose-cbor-encoded-cert-16, 25 January 2026,
              <https://datatracker.ietf.org/doc/html/draft-ietf-cose-
              cbor-encoded-cert-16>.

   [I-D.richardson-anima-masa-considerations]
              Richardson, M. and W. Pan, "Operational Considerations for
              Voucher infrastructure for BRSKI MASA", Work in Progress,
              Internet-Draft, draft-richardson-anima-masa-
              considerations-09, 22 January 2025,
              <https://datatracker.ietf.org/doc/html/draft-richardson-
              anima-masa-considerations-09>.

   [RFC4443]  Conta, A., Deering, S., and M. Gupta, Ed., "Internet
              Control Message Protocol (ICMPv6) for the Internet
              Protocol Version 6 (IPv6) Specification", STD 89,
              RFC 4443, DOI 10.17487/RFC4443, March 2006,
              <https://www.rfc-editor.org/rfc/rfc4443>.

   [RFC5652]  Housley, R., "Cryptographic Message Syntax (CMS)", STD 70,
              RFC 5652, DOI 10.17487/RFC5652, September 2009,
              <https://www.rfc-editor.org/rfc/rfc5652>.

   [RFC6282]  Hui, J., Ed. and P. Thubert, "Compression Format for IPv6
              Datagrams over IEEE 802.15.4-Based Networks", RFC 6282,
              DOI 10.17487/RFC6282, September 2011,
              <https://www.rfc-editor.org/rfc/rfc6282>.

   [RFC6775]  Shelby, Z., Ed., Chakrabarti, S., Nordmark, E., and C.
              Bormann, "Neighbor Discovery Optimization for IPv6 over
              Low-Power Wireless Personal Area Networks (6LoWPANs)",
              RFC 6775, DOI 10.17487/RFC6775, November 2012,
              <https://www.rfc-editor.org/rfc/rfc6775>.

   [RFC6838]  Freed, N., Klensin, J., and T. Hansen, "Media Type
              Specifications and Registration Procedures", BCP 13,
              RFC 6838, DOI 10.17487/RFC6838, January 2013,
              <https://www.rfc-editor.org/rfc/rfc6838>.

   [RFC7228]  Bormann, C., Ersue, M., and A. Keranen, "Terminology for
              Constrained-Node Networks", RFC 7228,
              DOI 10.17487/RFC7228, May 2014,
              <https://www.rfc-editor.org/rfc/rfc7228>.





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   [RFC7950]  Bjorklund, M., Ed., "The YANG 1.1 Data Modeling Language",
              RFC 7950, DOI 10.17487/RFC7950, August 2016,
              <https://www.rfc-editor.org/rfc/rfc7950>.

   [RFC8392]  Jones, M., Wahlstroem, E., Erdtman, S., and H. Tschofenig,
              "CBOR Web Token (CWT)", RFC 8392, DOI 10.17487/RFC8392,
              May 2018, <https://www.rfc-editor.org/rfc/rfc8392>.

   [RFC8990]  Bormann, C., Carpenter, B., Ed., and B. Liu, Ed., "GeneRic
              Autonomic Signaling Protocol (GRASP)", RFC 8990,
              DOI 10.17487/RFC8990, May 2021,
              <https://www.rfc-editor.org/rfc/rfc8990>.

   [RFC9334]  Birkholz, H., Thaler, D., Richardson, M., Smith, N., and
              W. Pan, "Remote ATtestation procedureS (RATS)
              Architecture", RFC 9334, DOI 10.17487/RFC9334, January
              2023, <https://www.rfc-editor.org/rfc/rfc9334>.

   [RFC9528]  Selander, G., Preuß Mattsson, J., and F. Palombini,
              "Ephemeral Diffie-Hellman Over COSE (EDHOC)", RFC 9528,
              DOI 10.17487/RFC9528, March 2024,
              <https://www.rfc-editor.org/rfc/rfc9528>.

   [RFC9595]  Veillette, M., Ed., Pelov, A., Ed., Petrov, I., Ed.,
              Bormann, C., and M. Richardson, "YANG Schema Item
              iDentifier (YANG SID)", RFC 9595, DOI 10.17487/RFC9595,
              July 2024, <https://www.rfc-editor.org/rfc/rfc9595>.

   [RFC9597]  Looker, T. and M.B. Jones, "CBOR Web Token (CWT) Claims in
              COSE Headers", RFC 9597, DOI 10.17487/RFC9597, June 2024,
              <https://www.rfc-editor.org/rfc/rfc9597>.

   [Thread]   Thread Group, Inc, "Thread Group website", February 2026,
              <https://www.threadgroup.org/>.

Appendix A.  Software and Library Support for cBRSKI

   This appendix lists software and security libraries that may be
   useful for implementing cBRSKI functionality.

A.1.  Open Source cBRSKI Implementations

   There are a few ongoing open source projects to support cBRSKI
   development and testing.  These include:

   *  OpenThread Registrar (OT Registrar) - a cBRSKI Registrar, test
      MASA server, and test Pledge written in Java.  Link
      (https://github.com/EskoDijk/ot-registrar)



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   *  OpenThread CCM (pre-alpha) - a cBRSKI Pledge and Join Proxy for
      OpenThread-based IoT nodes, written in C/C++. OpenThread nodes
      implement the [Thread] protocol.  Link
      (https://github.com/EskoDijk/openthread/pull/7)

   *  OpenThread Network Simulator v2 (OTNS2) - a CLI + GUI simulator
      for OpenThread IoT nodes in 6LoWPAN [RFC6282] mesh networks, able
      to accurately simulate cBRSKI Pledges onboarding (pre-alpha
      functionality) to a Thread mesh network via an OT Registrar.  Link
      (https://github.com/EskoDijk/ot-ns/pull/165)

   *  Fountain - a BRSKI/6TiSCH Registrar with support for COSE-signed
      vouchers, written in Ruby.  Link (https://github.com/AnimaGUS-
      minerva/fountain)

A.2.  Security Library Support

   For the implementation of BRSKI/cBRSKI, the use of a software library
   to manipulate PKIX certificates, establish secure (D)TLS connections,
   and use crypto algorithms is often beneficial.  Two C-based examples
   are OpenSSL and mbedtls.  Others more targeted to specific platforms
   or languages exist.  It is important to realize that the library
   interfaces differ significantly between libraries.

   Libraries do not support all known crypto algorithms.  Before
   deciding on a library, it is important to look at their supported
   crypto algorithms and the roadmap for future support.  Apart from
   availability, the library footprint, and the required execution
   cycles should be investigated beforehand.

   The handling of certificates usually includes the checking of a
   certificate chain.  In some libraries, chains are constructed and
   verified on the basis of a set of certificates, the trust anchor
   (usually a self signed root CA), and the target certificate.  In
   other libraries, the chain must be constructed beforehand and obey
   ordering criteria.  Verification always includes the checking of the
   signatures.  Less frequent is the checking the validity of the dates
   or checking the existence of a revoked certificate in the chain
   against a set of revoked certificates.  Checking the chain on the
   consistency of the certificate extensions which specify the use of
   the certificate usually needs to be programmed explicitly.

   A library can be used to construct a (D)TLS connection.  It is useful
   to realize that differences between (D)TLS implementations will occur
   due to the differences in the certificate checks supported by the
   library.  On top of that, checks between client and server
   certificates enforced by (D)TLS are not always helpful for a BRSKI
   implementation.  For example, the certificates of Pledge and



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   Registrar are usually not related when the BRSKI protocol is started.
   It must be verified that checks on the relation between client and
   server certificates do not hamper a succeful DTLS connection
   establishment.

A.2.1.  OpensSSL Example Code

   From OpenSSL's apps/verify.c :

   <CODE BEGINS>
   X509 *x = NULL;
   int i = 0, ret = 0;
   X509_STORE_CTX *csc;
   STACK_OF(X509) *chain = NULL;
   int num_untrusted;

   x = load_cert(file, "certificate file");
   if (x == NULL)
       goto end;

   csc = X509_STORE_CTX_new();
   if (csc == NULL) {
       BIO_printf(bio_err, "error %s: X.509 store context"
                  "allocation failed\n",
                  (file == NULL) ? "stdin" : file);
       goto end;
   }

   X509_STORE_set_flags(ctx, vflags);
   if (!X509_STORE_CTX_init(csc, ctx, x, uchain)) {
       X509_STORE_CTX_free(csc);
       BIO_printf(bio_err,
                  "error %s: X.509 store context"
                  "initialization failed\n",
                  (file == NULL) ? "stdin" : file);
       goto end;
   }
   if (tchain != NULL)
       X509_STORE_CTX_set0_trusted_stack(csc, tchain);
   if (crls != NULL)
       X509_STORE_CTX_set0_crls(csc, crls);

   i = X509_verify_cert(csc);
   X509_STORE_CTX_free(csc);

   <CODE ENDS>





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A.2.2.  mbedTLS Example Code

   <CODE BEGINS>
   mbedtls_x509_crt cert;
   mbedtls_x509_crt caCert;
   uint32_t         certVerifyResultFlags;
   // ...
   int result = mbedtls_x509_crt_verify(&cert, &caCert, NULL, NULL,
                                &certVerifyResultFlags, NULL, NULL);

   <CODE ENDS>

A.3.  Generating Certificates with OpenSSL

   This informative appendix shows example Bash shell scripts to
   generate test PKIX certificates for the Pledge IDevID, the Registrar
   and the MASA.  The shell scripts cannot be run stand-alone because
   they depend on input files which are not all included in this
   appendix.  Nevertheless, these scripts may provide guidance on how
   OpenSSL can be configured for generating cBRSKI certificates.

   The scripts were tested with OpenSSL 3.0.2.  Older versions may not
   work -- OpenSSL 1.1.1 for example does not support all extensions
   used.



























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   <CODE BEGINS>
   #!/bin/bash
   # File: create-cert-Pledge.sh
   # Create new cert for: Pledge IDevID

   # days certificate is valid - try to get close to the 802.1AR
   # specified 9999-12-31 end date.
   SECONDS1=`date +%s` # time now
   SECONDS2=`date --date="9999-12-31 23:59:59Z" +%s` # target end time
   let VALIDITY="(${SECONDS2}-${SECONDS1})/(24*3600)"
   echo "Using validity param -days ${VALIDITY}"

   NAME=pledge

   # create csr for device
   # conform to 802.1AR guidelines, using only CN + serialNumber when
   # manufacturer is already present as CA.
   # CN is not even mandatory, but just good practice.
   openssl req -new -key keys/privkey_pledge.pem -out $NAME.csr -subj \
                "/CN=Stok IoT sensor Y-42/serialNumber=JADA123456789"

   # sign csr - it uses faketime only to get endtime to 23:59:59Z
   faketime '23:59:59Z' \
   openssl x509 -set_serial 32429 -CAform PEM -CA output/masa_ca.pem \
     -CAkey keys/privkey_masa_ca.pem -extfile x509v3.ext -extensions \
     pledge_ext -req -in $NAME.csr -out output/$NAME.pem \
     -days $VALIDITY -sha256

   # Note: alternative method using 'ca' command. Currently
   # doesn't work without 'country' subject field.
   # openssl ca -rand_serial -enddate 99991231235959Z -certform PEM \
   #  -cert output/masa_ca.pem -keyfile keys/privkey_masa_ca.pem \
   #  -extfile x509v3.ext -extensions pledge_ext -in $NAME.csr \
   #  -out $NAME.pem -outdir output

   # delete temp files
   rm -f $NAME.csr

   # convert to .der format
   openssl x509 -in output/$NAME.pem -inform PEM -out output/$NAME.der \
                -outform DER

   <CODE ENDS>








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   <CODE BEGINS>
   # File: x509v3.ext
   # This file contains all X509v3 extension definitions for OpenSSL
   # certificate generation. Each certificate has its own _ext
   # section below.

   [ req ]
   prompt = no

   [ masa_ca_ext ]
   subjectAltName=email:info@masa.stok.nl
   keyUsage = critical,digitalSignature, keyCertSign, cRLSign
   basicConstraints = critical,CA:TRUE,pathlen:3
   subjectKeyIdentifier=hash
   authorityKeyIdentifier=keyid

   [ pledge_ext ]
   keyUsage = critical,digitalSignature, nonRepudiation, \
              keyEncipherment, dataEncipherment
   # basicConstraints for a non-CA cert MAY be marked either
   # non-critical or critical.
   basicConstraints = CA:FALSE
   # Don't include subjectKeyIdentifier (SKI) - see 802.1AR-2018
   subjectKeyIdentifier = none
   authorityKeyIdentifier=keyid
   # Include the MASA URI
   1.3.6.1.5.5.7.1.32 = ASN1:IA5STRING:masa.stok.nl

   [ domain_ca_ext ]
   subjectAltName=email:help@custom-er.example.com
   keyUsage = critical, keyCertSign, digitalSignature, cRLSign
   basicConstraints=critical,CA:TRUE
   # RFC 5280 4.2.1.1 : AKI MAY be omitted, and MUST be non-critical;
   # SKI MUST be non-critical
   subjectKeyIdentifier=hash

   [ registrar_ext ]
   keyUsage = critical, digitalSignature, nonRepudiation, \
              keyEncipherment, dataEncipherment
   basicConstraints=CA:FALSE
   subjectKeyIdentifier=hash
   authorityKeyIdentifier=keyid
   # Set Registrar 'RA' flag along with TLS client/server usage
   #  see draft-ietf-anima-constrained-voucher#section-7.3
   #  see tools.ietf.org/html/rfc6402#section-2.10
   #  see www.openssl.org/docs/man1.1.1/man5/x509v3_config.html
   extendedKeyUsage = critical,1.3.6.1.5.5.7.3.28, serverAuth, \
                      clientAuth



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   <CODE ENDS>

   <CODE BEGINS>
   #!/bin/bash
   # File: create-cert-Registrar.sh
   # Create new cert for: Registrar in a company domain

   # days certificate is valid
   VALIDITY=1095

   # cert filename
   NAME=registrar

   # create csr
   openssl req -new -key keys/privkey_registrar.pem -out $NAME.csr \
    -subj "/CN=Custom-ER Registrar/OU=Office dept/O=Custom-ER, Inc./\
   L=Ottowa/ST=ON/C=CA"

   # sign csr
   openssl x509 -set_serial 0xC3F62149B2E30E3E -CAform PEM -CA \
    output/domain_ca.pem -extfile x509v3.ext -extensions registrar_ext \
    -req -in $NAME.csr -CAkey keys/privkey_domain_ca.pem \
    -out output/$NAME.pem -days $VALIDITY -sha256

   # delete temp files
   rm -f $NAME.csr

   # convert to .der format
   openssl x509 -in output/$NAME.pem -inform PEM -out output/$NAME.der \
                -outform DER

   <CODE ENDS>



















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   <CODE BEGINS>
   #!/bin/bash
   # File: create-cert-MASA.sh
   # Create new cert for: MASA CA, self-signed CA certificate

   # days certificate is valid
   VALIDITY=3650

   NAME=masa_ca

   # create csr
   openssl req -new -key keys/privkey_masa_ca.pem -out $NAME.csr \
               -subj "/CN=masa.stok.nl/O=vanderstok/L=Helmond/C=NL"

   # sign csr
   mkdir output >& /dev/null
   openssl x509 -set_serial 0xE39CDA17E1386A0A  -extfile x509v3.ext \
    -extensions masa_ca_ext -req -in $NAME.csr \
    -signkey keys/privkey_masa_ca.pem -out output/$NAME.pem \
    -days $VALIDITY -sha256

   # delete temp files
   rm -f $NAME.csr

   # convert to .der format
   openssl x509 -in output/$NAME.pem -inform PEM -out output/$NAME.der \
                -outform DER

   <CODE ENDS>

Appendix B.  cBRSKI Message Examples

   This appendix extends the EST-coaps message examples from Appendix A
   of [RFC9148] with cBRSKI messages.  The CoAP headers are only fully
   worked out for the first example, enrollstatus.

B.1.  enrollstatus

   A coaps enrollstatus message from Pledge to Registrar can be as
   follows:

     REQ: POST /b/es
       Content-Format: 60 (application/cbor)
       Payload: <binary CBOR encoding of an enrollstatus map>

   The corresponding CoAP header fields for this request are shown
   below.




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     Ver = 1
     T = 0 (CON)
     TKL = 1
     Code = 0x02 (0.02 is POST method)
     Message ID = 0xab0f
     Token = 0x4d
     Options
      Option  (Uri-Path)
        Option Delta = 0xb   (option nr = 11)
        Option Length = 0x1
        Option Value = "b"
      Option  (Uri-Path)
        Option Delta = 0x0   (option nr = 11)
        Option Length = 0x2
        Option Value = "es"
      Option  (Content-Format)
        Option Delta = 0x1   (option nr = 12)
        Option Length = 0x1
        Option Value = 60    (application/cbor)
     Payload Marker = 0xFF
     Payload = A26776657273696F6E0166737461747573F5 (18 bytes binary)

   The Uri-Host and Uri-Port Options are omitted because they coincide
   with the transport protocol (UDP) destination address and port
   respectively.

   The above binary CBOR enrollstatus payload looks as follows in CBOR
   diagnostic notation, for the case of enrollment success:

     {
       "version": 1,
       "status": true
      }

   Alternatively the payload could look as follows in case of enrollment
   failure, using the 'reason' map item value to describe the failure:

     Payload = A36776657273696F6E0166737461747573F466726561736F6E782A3C
               496E666F726D61746976652068756D616E207265616461626C652065
               72726F72206D6573736167653E    (69 bytes binary)

     {
       "version": 1,
       "status": false,
       "reason": "<Informative human readable error message>"
     }





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   To indicate successful reception of the enrollmentstatus telemetry
   report, a response from the Registrar may then be:

     2.04 Changed

   Which in case of a piggybacked response has the following CoAP header
   fields:

     Ver=1
     T=2 (ACK)
     TKL=1
     Code = 0x44 (2.04 Changed)
     Message ID = 0xab0f
     Token = 0x4d

B.2.  voucher_status

   A coaps voucher_status message from Pledge to Registrar can be as
   follows:

     REQ: POST /.well-known/brski/vs
       Content-Format: 60 (application/cbor)
       Payload:
       A46776657273696F6E0166737461747573F466726561736F6E7828496E66
       6F726D61746976652068756D616E2D7265616461626C65206572726F7220
       6D6573736167656E726561736F6E2D636F6E74657874A100764164646974
       696F6E616C20696E666F726D6174696F6E

   The request payload above is binary CBOR but represented here in
   hexadecimal for readability.  Below is the equivalent CBOR diagnostic
   format.

     {
       "version": 1,
       "status": false,
       "reason": "Informative human-readable error message",
       "reason-context": { 0: "Additional information" }
     }

   A success response without payload will then be sent by the Registrar
   back to the Pledge to indicate reception of the telemetry report:

     RES: 2.04 Changed








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Appendix C.  COSE-signed Voucher (Request) Examples

   This appendix provides examples of COSE-signed voucher requests and
   vouchers.  First, the used test keys and PKIX certificates are
   described, followed by examples of a constrained PVR, RVR and
   voucher.

C.1.  Pledge, Registrar and MASA Keys

   This section documents the public and private keys used for all
   examples in this appendix.  These keys are not used in any production
   system, and must only be used for testing purposes.

C.1.1.  Pledge IDevID Private Key

   -----BEGIN EC PRIVATE KEY-----
   MHcCAQEEIMv+C4dbzeyrEH20qkpFlWIH2FFACGZv9kW7rNWtSlYtoAoGCCqGSM49
   AwEHoUQDQgAESH6OUiYFRhfIgWl4GG8jHoj8a+8rf6t5s1mZ/4SePlKom39GQ34p
   VYryJ9aHmboLLfz69bzICQFKbkoQ5oaiew==
   -----END EC PRIVATE KEY-----

   Private-Key: (256 bit)
   priv:
       cb:fe:0b:87:5b:cd:ec:ab:10:7d:b4:aa:4a:45:95:
       62:07:d8:51:40:08:66:6f:f6:45:bb:ac:d5:ad:4a:
       56:2d
   pub:
       04:48:7e:8e:52:26:05:46:17:c8:81:69:78:18:6f:
       23:1e:88:fc:6b:ef:2b:7f:ab:79:b3:59:99:ff:84:
       9e:3e:52:a8:9b:7f:46:43:7e:29:55:8a:f2:27:d6:
       87:99:ba:0b:2d:fc:fa:f5:bc:c8:09:01:4a:6e:4a:
       10:e6:86:a2:7b
   ASN1 OID: prime256v1
   NIST CURVE: P-256

C.1.2.  Registrar Private Key

   -----BEGIN PRIVATE KEY-----
   MIGHAgEAMBMGByqGSM49AgEGCCqGSM49AwEHBG0wawIBAQQgYJ/MP0dWA9BkYd4W
   s6oRY62hDddaEmrAVm5dtAXE/UGhRANCAAQgMIVb6EaRCz7LFcr4Vy0+tWW9xlSh
   Xvr27euqi54WCMXJEMk6IIaPyFBNNw8bJvqXWfZ5g7t4hj7amsvqUST2
   -----END PRIVATE KEY-----









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   Private-Key: (256 bit)
   priv:
       60:9f:cc:3f:47:56:03:d0:64:61:de:16:b3:aa:11:
       63:ad:a1:0d:d7:5a:12:6a:c0:56:6e:5d:b4:05:c4:
       fd:41
   pub:
       04:20:30:85:5b:e8:46:91:0b:3e:cb:15:ca:f8:57:
       2d:3e:b5:65:bd:c6:54:a1:5e:fa:f6:ed:eb:aa:8b:
       9e:16:08:c5:c9:10:c9:3a:20:86:8f:c8:50:4d:37:
       0f:1b:26:fa:97:59:f6:79:83:bb:78:86:3e:da:9a:
       cb:ea:51:24:f6
   ASN1 OID: prime256v1
   NIST CURVE: P-256

C.1.3.  MASA Private Key

   -----BEGIN PRIVATE KEY-----
   MIGHAgEAMBMGByqGSM49AgEGCCqGSM49AwEHBG0wawIBAQQgrbJ1oU+HIJ2SWYAk
   DkBTL+YNPxQG+gwsMsZB94N8mZ2hRANCAASS9NVlWJdztwNY81yPlH2UODYWhlYA
   ZfsqnEPSFZKnq8mq8gF78ZVbYi6q2FEg8kkORY/rpIU/X7SQsRuD+wMW
   -----END PRIVATE KEY-----

   Private-Key: (256 bit)
   priv:
       ad:b2:75:a1:4f:87:20:9d:92:59:80:24:0e:40:53:
       2f:e6:0d:3f:14:06:fa:0c:2c:32:c6:41:f7:83:7c:
       99:9d
   pub:
       04:92:f4:d5:65:58:97:73:b7:03:58:f3:5c:8f:94:
       7d:94:38:36:16:86:56:00:65:fb:2a:9c:43:d2:15:
       92:a7:ab:c9:aa:f2:01:7b:f1:95:5b:62:2e:aa:d8:
       51:20:f2:49:0e:45:8f:eb:a4:85:3f:5f:b4:90:b1:
       1b:83:fb:03:16
   ASN1 OID: prime256v1
   NIST CURVE: P-256

C.2.  Pledge, Registrar, Domain CA and MASA Certificates

   All keys and PKIX certificates used for the examples have been
   generated with OpenSSL - see Appendix A.3 for more details on
   certificate generation.  Below the certificates are listed that
   accompany the keys shown above.  Each certificate description is
   followed by the hexadecimal representation of the X.509 ASN.1 DER
   encoded certificate.  This representation can be for example decoded
   using an online ASN.1 decoder.

C.2.1.  Pledge IDevID Certificate




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   Certificate:
   Data:
    Version: 3 (0x2)
    Serial Number: 32429 (0x7ead)
    Signature Algorithm: ecdsa-with-SHA256
    Issuer: CN = masa.stok.nl, O = vanderstok, L = Helmond,
            C = NL
    Validity
      Not Before: Dec  9 12:50:47 2022 GMT
      Not After : Dec 31 12:50:47 9999 GMT
    Subject: CN = Stok IoT sensor Y-42, serialNumber = JADA123456789
    Subject Public Key Info:
      Public Key Algorithm: id-ecPublicKey
        Public-Key: (256 bit)
        pub:
          04:48:7e:8e:52:26:05:46:17:c8:81:69:78:18:6f:
          23:1e:88:fc:6b:ef:2b:7f:ab:79:b3:59:99:ff:84:
          9e:3e:52:a8:9b:7f:46:43:7e:29:55:8a:f2:27:d6:
          87:99:ba:0b:2d:fc:fa:f5:bc:c8:09:01:4a:6e:4a:
          10:e6:86:a2:7b
        ASN1 OID: prime256v1
        NIST CURVE: P-256
    X509v3 extensions:
      X509v3 Key Usage: critical
        Digital Signature, Non Repudiation, Key Encipherment,
                Data Encipherment
      X509v3 Basic Constraints:
        CA:FALSE
      X509v3 Authority Key Identifier:
        CB:8D:98:CA:74:C5:1B:58:DD:E7:AC:EF:86:9A:94:43:A8:D6:66:A6
      1.3.6.1.5.5.7.1.32:
         hl=2 l=  12 prim: IA5STRING     :masa.stok.nl

   Signature Algorithm: ecdsa-with-SHA256
   Signature Value:
    30:45:02:20:4d:89:90:7e:03:fb:52:56:42:0c:3f:c1:b1:f1:
    47:b5:b3:93:65:45:2e:be:50:db:67:85:8f:23:89:a2:3f:9e:
    02:21:00:95:33:69:d1:c6:db:f0:f1:f6:52:24:59:d3:0a:95:
    4e:b2:f4:96:a1:31:3c:7b:d9:2f:28:b3:29:71:bb:60:df

   Below is the hexadecimal representation of the binary X.509 DER-
   encoded certificate:









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   308201CE30820174A00302010202027EAD300A06082A8648CE3D040302304B31
   15301306035504030C0C6D6173612E73746F6B2E6E6C31133011060355040A0C
   0A76616E64657273746F6B3110300E06035504070C0748656C6D6F6E64310B30
   09060355040613024E4C3020170D3232313230393132353034375A180F393939
   39313233313132353034375A3037311D301B06035504030C1453746F6B20496F
   542073656E736F7220592D3432311630140603550405130D4A41444131323334
   35363738393059301306072A8648CE3D020106082A8648CE3D03010703420004
   487E8E5226054617C8816978186F231E88FC6BEF2B7FAB79B35999FF849E3E52
   A89B7F46437E29558AF227D68799BA0B2DFCFAF5BCC809014A6E4A10E686A27B
   A35A3058300E0603551D0F0101FF0404030204F030090603551D130402300030
   1F0603551D23041830168014CB8D98CA74C51B58DDE7ACEF869A9443A8D666A6
   301A06082B06010505070120040E160C6D6173612E73746F6B2E6E6C300A0608
   2A8648CE3D040302034800304502204D89907E03FB5256420C3FC1B1F147B5B3
   9365452EBE50DB67858F2389A23F9E022100953369D1C6DBF0F1F6522459D30A
   954EB2F496A1313C7BD92F28B32971BB60DF

C.2.2.  Registrar Certificate


































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   Certificate:
   Data:
    Version: 3 (0x2)
    Serial Number:
      c3:f6:21:49:b2:e3:0e:3e
    Signature Algorithm: ecdsa-with-SHA256
    Issuer: CN = Custom-ER Global CA, OU = IT, O = "Custom-ER, Inc.",
            L = San Jose, ST = CA, C = US
    Validity
      Not Before: Dec  9 12:50:47 2022 GMT
      Not After : Dec  8 12:50:47 2025 GMT
    Subject: CN = Custom-ER Registrar, OU = Office dept, O = "Custom-ER,
            Inc.", L = Ottowa, ST = ON, C = CA
    Subject Public Key Info:
      Public Key Algorithm: id-ecPublicKey
        Public-Key: (256 bit)
        pub:
          04:20:30:85:5b:e8:46:91:0b:3e:cb:15:ca:f8:57:
          2d:3e:b5:65:bd:c6:54:a1:5e:fa:f6:ed:eb:aa:8b:
          9e:16:08:c5:c9:10:c9:3a:20:86:8f:c8:50:4d:37:
          0f:1b:26:fa:97:59:f6:79:83:bb:78:86:3e:da:9a:
          cb:ea:51:24:f6
        ASN1 OID: prime256v1
        NIST CURVE: P-256
    X509v3 extensions:
      X509v3 Key Usage: critical
        Digital Signature, Non Repudiation, Key Encipherment,
                Data Encipherment
      X509v3 Basic Constraints:
        CA:FALSE
      X509v3 Subject Key Identifier:
        C9:08:0B:38:7D:8D:D8:5B:3A:59:E7:EC:10:0B:86:63:93:A9:CA:4C
      X509v3 Authority Key Identifier:
        92:EA:76:40:40:4A:8F:AB:4F:27:0B:F3:BC:37:9D:86:CD:72:80:F8
      X509v3 Extended Key Usage: critical
        CMC Registration Authority, TLS Web Server Authentication,
                TLS Web Client Authentication
   Signature Algorithm: ecdsa-with-SHA256
   Signature Value:
    30:45:02:21:00:d8:4a:7c:69:2f:f9:58:6e:82:22:87:18:f6:
    3b:c3:05:f0:ae:b8:ae:ec:42:78:82:38:79:81:2a:5d:15:61:
    64:02:20:08:f2:3c:13:69:13:b0:2c:e2:63:09:d5:99:4f:eb:
    75:70:af:af:ed:98:cd:f1:12:11:c0:37:f7:18:4d:c1:9d

   Below is the hexadecimal representation of the binary X.509 DER-
   encoded certificate:





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   3082026D30820213A003020102020900C3F62149B2E30E3E300A06082A8648CE
   3D0403023072311C301A06035504030C13437573746F6D2D455220476C6F6261
   6C204341310B3009060355040B0C02495431183016060355040A0C0F43757374
   6F6D2D45522C20496E632E3111300F06035504070C0853616E204A6F7365310B
   300906035504080C024341310B3009060355040613025553301E170D32323132
   30393132353034375A170D3235313230383132353034375A3079311C301A0603
   5504030C13437573746F6D2D4552205265676973747261723114301206035504
   0B0C0B4F6666696365206465707431183016060355040A0C0F437573746F6D2D
   45522C20496E632E310F300D06035504070C064F74746F7761310B3009060355
   04080C024F4E310B30090603550406130243413059301306072A8648CE3D0201
   06082A8648CE3D030107034200042030855BE846910B3ECB15CAF8572D3EB565
   BDC654A15EFAF6EDEBAA8B9E1608C5C910C93A20868FC8504D370F1B26FA9759
   F67983BB78863EDA9ACBEA5124F6A3818A308187300E0603551D0F0101FF0404
   030204F030090603551D1304023000301D0603551D0E04160414C9080B387D8D
   D85B3A59E7EC100B866393A9CA4C301F0603551D2304183016801492EA764040
   4A8FAB4F270BF3BC379D86CD7280F8302A0603551D250101FF0420301E06082B
   0601050507031C06082B0601050507030106082B06010505070302300A06082A
   8648CE3D0403020348003045022100D84A7C692FF9586E82228718F63BC305F0
   AEB8AEEC4278823879812A5D156164022008F23C136913B02CE26309D5994FEB
   7570AFAFED98CDF11211C037F7184DC19D

C.2.3.  Domain CA Certificate

   The Domain CA certificate is the CA of the owner's domain.  It has
   signed the Registrar (RA) certificate.


























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   Certificate:
   Data:
    Version: 3 (0x2)
    Serial Number: 3092288576548618702 (0x2aea0413a42dc1ce)
    Signature Algorithm: ecdsa-with-SHA256
    Issuer: CN = Custom-ER Global CA, OU = IT, O = "Custom-ER, Inc.",
            L = San Jose, ST = CA, C = US
    Validity
      Not Before: Dec  9 12:50:47 2022 GMT
      Not After : Dec  6 12:50:47 2032 GMT
    Subject: CN = Custom-ER Global CA, OU = IT, O = "Custom-ER, Inc.",
            L = San Jose, ST = CA, C = US
    Subject Public Key Info:
      Public Key Algorithm: id-ecPublicKey
        Public-Key: (256 bit)
        pub:
          04:97:b1:ed:96:91:64:93:09:85:bb:b8:ac:9a:2a:
          f9:45:5c:df:ee:a4:b1:1d:e2:e7:9d:06:8b:fa:80:
          39:26:b4:00:52:51:b3:4f:1c:08:15:a4:cb:e0:3f:
          bd:1b:bc:b6:35:f6:43:1a:22:de:78:65:3b:87:b9:
          95:37:ec:e1:6c
        ASN1 OID: prime256v1
        NIST CURVE: P-256
    X509v3 extensions:
      X509v3 Subject Alternative Name:
        email:help@custom-er.example.com
      X509v3 Key Usage: critical
        Digital Signature, Certificate Sign, CRL Sign
      X509v3 Basic Constraints: critical
        CA:TRUE
      X509v3 Subject Key Identifier:
        92:EA:76:40:40:4A:8F:AB:4F:27:0B:F3:BC:37:9D:86:CD:72:80:F8
   Signature Algorithm: ecdsa-with-SHA256
   Signature Value:
    30:44:02:20:66:15:df:c3:70:11:f6:73:78:d8:fd:1c:2a:3f:
    bd:d1:3f:51:f6:b6:6f:2d:7c:e2:7a:13:18:21:bb:70:f0:c0:
    02:20:69:86:d8:d2:28:b2:92:6e:23:9e:19:0b:8f:18:25:c9:
    c1:4c:67:95:ff:a0:b3:24:bd:4d:ac:2e:cb:68:d7:13

   Below is the hexadecimal representation of the binary X.509 DER-
   encoded certificate:










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   30820242308201E9A00302010202082AEA0413A42DC1CE300A06082A8648CE3D
   0403023072311C301A06035504030C13437573746F6D2D455220476C6F62616C
   204341310B3009060355040B0C02495431183016060355040A0C0F437573746F
   6D2D45522C20496E632E3111300F06035504070C0853616E204A6F7365310B30
   0906035504080C024341310B3009060355040613025553301E170D3232313230
   393132353034375A170D3332313230363132353034375A3072311C301A060355
   04030C13437573746F6D2D455220476C6F62616C204341310B3009060355040B
   0C02495431183016060355040A0C0F437573746F6D2D45522C20496E632E3111
   300F06035504070C0853616E204A6F7365310B300906035504080C024341310B
   30090603550406130255533059301306072A8648CE3D020106082A8648CE3D03
   01070342000497B1ED969164930985BBB8AC9A2AF9455CDFEEA4B11DE2E79D06
   8BFA803926B4005251B34F1C0815A4CBE03FBD1BBCB635F6431A22DE78653B87
   B99537ECE16CA369306730250603551D11041E301C811A68656C704063757374
   6F6D2D65722E6578616D706C652E636F6D300E0603551D0F0101FF0404030201
   86300F0603551D130101FF040530030101FF301D0603551D0E0416041492EA76
   40404A8FAB4F270BF3BC379D86CD7280F8300A06082A8648CE3D040302034700
   304402206615DFC37011F67378D8FD1C2A3FBDD13F51F6B66F2D7CE27A131821
   BB70F0C002206986D8D228B2926E239E190B8F1825C9C14C6795FFA0B324BD4D
   AC2ECB68D713

C.2.4.  MASA Certificate

   The MASA CA certificate is the CA that signed the Pledge's IDevID
   certificate.



























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   Certificate:
   Data:
    Version: 3 (0x2)
    Serial Number:
      e3:9c:da:17:e1:38:6a:0a
    Signature Algorithm: ecdsa-with-SHA256
    Issuer: CN = masa.stok.nl, O = vanderstok, L = Helmond,
            C = NL
    Validity
      Not Before: Dec  9 12:50:47 2022 GMT
      Not After : Dec  6 12:50:47 2032 GMT
    Subject: CN = masa.stok.nl, O = vanderstok, L = Helmond,
            C = NL
    Subject Public Key Info:
      Public Key Algorithm: id-ecPublicKey
        Public-Key: (256 bit)
        pub:
          04:92:f4:d5:65:58:97:73:b7:03:58:f3:5c:8f:94:
          7d:94:38:36:16:86:56:00:65:fb:2a:9c:43:d2:15:
          92:a7:ab:c9:aa:f2:01:7b:f1:95:5b:62:2e:aa:d8:
          51:20:f2:49:0e:45:8f:eb:a4:85:3f:5f:b4:90:b1:
          1b:83:fb:03:16
        ASN1 OID: prime256v1
        NIST CURVE: P-256
    X509v3 extensions:
      X509v3 Subject Alternative Name:
        email:info@masa.stok.nl
      X509v3 Key Usage: critical
        Digital Signature, Certificate Sign, CRL Sign
      X509v3 Basic Constraints: critical
        CA:TRUE, pathlen:3
      X509v3 Subject Key Identifier:
        CB:8D:98:CA:74:C5:1B:58:DD:E7:AC:EF:86:9A:94:43:A8:D6:66:A6
   Signature Algorithm: ecdsa-with-SHA256
   Signature Value:
    30:46:02:21:00:94:3f:a5:26:51:68:16:38:5b:78:9a:d8:c3:
    af:8e:49:28:22:60:56:26:43:4a:14:98:3e:e1:e4:81:ad:ca:
    1b:02:21:00:ba:4d:aa:fd:fa:68:42:74:03:2b:a8:41:6b:e2:
    90:0c:9e:7b:b8:c0:9c:f7:0e:3f:b4:36:8a:b3:9c:3e:31:0e

   Below is the hexadecimal representation of the binary X.509 DER-
   encoded certificate:









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   308201F130820196A003020102020900E39CDA17E1386A0A300A06082A8648CE
   3D040302304B3115301306035504030C0C6D6173612E73746F6B2E6E6C311330
   11060355040A0C0A76616E64657273746F6B3110300E06035504070C0748656C
   6D6F6E64310B3009060355040613024E4C301E170D3232313230393132353034
   375A170D3332313230363132353034375A304B3115301306035504030C0C6D61
   73612E73746F6B2E6E6C31133011060355040A0C0A76616E64657273746F6B31
   10300E06035504070C0748656C6D6F6E64310B3009060355040613024E4C3059
   301306072A8648CE3D020106082A8648CE3D0301070342000492F4D565589773
   B70358F35C8F947D9438361686560065FB2A9C43D21592A7ABC9AAF2017BF195
   5B622EAAD85120F2490E458FEBA4853F5FB490B11B83FB0316A3633061301C06
   03551D11041530138111696E666F406D6173612E73746F6B2E6E6C300E060355
   1D0F0101FF04040302018630120603551D130101FF040830060101FF02010330
   1D0603551D0E04160414CB8D98CA74C51B58DDE7ACEF869A9443A8D666A6300A
   06082A8648CE3D0403020349003046022100943FA526516816385B789AD8C3AF
   8E492822605626434A14983EE1E481ADCA1B022100BA4DAAFDFA684274032BA8
   416BE2900C9E7BB8C09CF70E3FB4368AB39C3E310E

C.3.  COSE-signed Pledge Voucher Request (PVR)

   In this example, the voucher request (PVR) has been signed by the
   Pledge using the IDevID private key of Appendix C.1.1, and has been
   sent to the link-local constrained Join Proxy (JP) over CoAPS to JP's
   join port.  The join port happens to use the default CoAPS UDP port
   5684.

     REQ: POST coaps://[JP-link-local-address]/b/rv
       Content-Format: 836 (application/voucher+cose)
       Payload: <signed_pvr>

   When the Join Proxy receives the DTLS handshake messages from the
   Pledge, it will relay these messages to the Registrar.  The payload
   signed_voucher_request is shown as hexadecimal dump (with lf added)
   below:

   D28443A10126A0587EA11909C5A40102074823BFBBC9C2BCF2130C585B305930
   1306072A8648CE3D020106082A8648CE3D030107034200042030855BE846910B
   3ECB15CAF8572D3EB565BDC654A15EFAF6EDEBAA8B9E1608C5C910C93A20868F
   C8504D370F1B26FA9759F67983BB78863EDA9ACBEA5124F60D6D4A4144413132
   33343536373839584068987DE8B007F4E9416610BBE2D48E1D7EA1032092B8BF
   CE611421950F45B22F17E214820C07E777ADF86175E25D3205568404C25FCEEC
   1B817C7861A6104B3D

   The representation of signed_pvr in CBOR diagnostic format (with lf
   added) is:







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   18([h'A10126', {}, h'A11909C5A40102074823BFBBC9C2BCF2130C585B3059301
   306072A8648CE3D020106082A8648CE3D030107034200042030855BE846910B3ECB1
   5CAF8572D3EB565BDC654A15EFAF6EDEBAA8B9E1608C5C910C93A20868FC8504D370
   F1B26FA9759F67983BB78863EDA9ACBEA5124F60D6D4A41444131323334353637383
   9', h'68987DE8B007F4E9416610BBE2D48E1D7EA1032092B8BFCE611421950F45B2
   2F17E214820C07E777ADF86175E25D3205568404C25FCEEC1B817C7861A6104B3D']
   )

   The COSE payload is the PVR voucher data, encoded as a CBOR byte
   string.  The diagnostic representation of it is shown below:

   {2501: {1: 2, 7: h'23BFBBC9C2BCF213', 12: h'3059301306072A8648CE3D02
   0106082A8648CE3D030107034200042030855BE846910B3ECB15CAF8572D3EB565BD
   C654A15EFAF6EDEBAA8B9E1608C5C910C93A20868FC8504D370F1B26FA9759F67983
   BB78863EDA9ACBEA5124F6', 13: "JADA123456789"}}

   The Pledge uses the 'proximity' (key '1', SID 2502, enum value 2)
   assertion together with an included 'proximity-registrar-pubk'
   attribute (key '12', SID 2513) to inform MASA about its proximity to
   the specific Registrar.

C.4.  COSE-signed Registrar Voucher Request (RVR)

   In this example the Registrar's voucher request has been signed by
   the JRC (Registrar) using the private key from Appendix C.1.2.
   Contained within this voucher request is the voucher request PVR that
   was made by the Pledge to JRC.  Note that the RVR uses the HTTPS
   protocol (not CoAP) and corresponding long URI path names as defined
   in [RFC8995].  The Content-Type and Accept headers indicate the
   constrained voucher format that is defined in the present document.
   Because the Pledge used this format in the PVR, the JRC must also use
   this format in the RVR.

     REQ: POST https://masa.stok.nl/.well-known/brski/requestvoucher
       Content-Type: application/voucher+cose
       Accept: application/voucher+cose
       Body: <signed_rvr>

   The payload signed_rvr is shown as hexadecimal dump (with lf added):

   D28443A10126A11820825902843082028030820225A003020102020900C3F621
   49B2E30E3E300A06082A8648CE3D0403023072311C301A06035504030C134375
   73746F6D2D455220476C6F62616C204341310B3009060355040B0C0249543118
   3016060355040A0C0F437573746F6D2D45522C20496E632E3111300F06035504
   070C0853616E204A6F7365310B300906035504080C024341310B300906035504
   0613025553301E170D3232313230363131333735395A170D3235313230353131
   333735395A30818D3131302F06035504030C28437573746F6D2D455220436F6D
   6D65726369616C204275696C64696E6773205265676973747261723113301106



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   0355040B0C0A4F6666696365206F707331183016060355040A0C0F437573746F
   6D2D45522C20496E632E310F300D06035504070C064F74746F7761310B300906
   035504080C024F4E310B30090603550406130243413059301306072A8648CE3D
   020106082A8648CE3D030107034200042030855BE846910B3ECB15CAF8572D3E
   B565BDC654A15EFAF6EDEBAA8B9E1608C5C910C93A20868FC8504D370F1B26FA
   9759F67983BB78863EDA9ACBEA5124F6A3818730818430090603551D13040230
   00300B0603551D0F0404030204F0301D0603551D0E04160414C9080B387D8DD8
   5B3A59E7EC100B866393A9CA4C301F0603551D2304183016801492EA7640404A
   8FAB4F270BF3BC379D86CD7280F8302A0603551D250101FF0420301E06082B06
   01050507031C06082B0601050507030106082B06010505070302300A06082A86
   48CE3D040302034900304602210091A2033692EB81503D53505FFC8DA326B1EE
   7DEA96F29174F0B3341A07812201022100FF7339288108B712F418530A18025A
   895408CC45E0BB678B46FBAB37DDB4D36B59024730820243308201E9A0030201
   0202082AEA0413A42DC1CE300A06082A8648CE3D0403023072311C301A060355
   04030C13437573746F6D2D455220476C6F62616C204341310B3009060355040B
   0C02495431183016060355040A0C0F437573746F6D2D45522C20496E632E3111
   300F06035504070C0853616E204A6F7365310B300906035504080C024341310B
   3009060355040613025553301E170D3232313230363131333735395A170D3332
   313230333131333735395A3072311C301A06035504030C13437573746F6D2D45
   5220476C6F62616C204341310B3009060355040B0C0249543118301606035504
   0A0C0F437573746F6D2D45522C20496E632E3111300F06035504070C0853616E
   204A6F7365310B300906035504080C024341310B300906035504061302555330
   59301306072A8648CE3D020106082A8648CE3D0301070342000497B1ED969164
   930985BBB8AC9A2AF9455CDFEEA4B11DE2E79D068BFA803926B4005251B34F1C
   0815A4CBE03FBD1BBCB635F6431A22DE78653B87B99537ECE16CA3693067300F
   0603551D130101FF040530030101FF30250603551D11041E301C811A68656C70
   40637573746F6D2D65722E6578616D706C652E636F6D300E0603551D0F0101FF
   040403020186301D0603551D0E0416041492EA7640404A8FAB4F270BF3BC379D
   86CD7280F8300A06082A8648CE3D0403020348003045022100D6D813B390BD3A
   7B4E85424BCB1ED933AD1E981F2817B59083DD6EC1C5E3FADF02202CEE440619
   2BC767E98D7CFAE044C6807481AD8564A7D569DCA3D1CDF1E5E843590124A119
   09C5A60102027818323032322D31322D30365432303A30343A31352E3735345A
   05581A041830168014CB8D98CA74C51B58DDE7ACEF869A9443A8D666A6074823
   BFBBC9C2BCF2130958C9D28443A10126A0587EA11909C5A40102074823BFBBC9
   C2BCF2130C585B3059301306072A8648CE3D020106082A8648CE3D0301070342
   00042030855BE846910B3ECB15CAF8572D3EB565BDC654A15EFAF6EDEBAA8B9E
   1608C5C910C93A20868FC8504D370F1B26FA9759F67983BB78863EDA9ACBEA51
   24F60D6D4A414441313233343536373839584068987DE8B007F4E9416610BBE2
   D48E1D7EA1032092B8BFCE611421950F45B22F17E214820C07E777ADF86175E2
   5D3205568404C25FCEEC1B817C7861A6104B3D0D6D4A41444131323334353637
   38395840B1DD40B10787437588AEAC9036899191C16CCDBECA31C197855CCB6B
   BA142D709FE329CBC3F76297D6063ACB6759EAB98E96EA4C4AA2135AA48A247B
   AC1D6A3F

   The representation of signed_rvr in CBOR diagnostic format (with lf
   added) is:





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   18([h'A10126', {32: [h'3082028030820225A003020102020900C3F62149B2E30
   E3E300A06082A8648CE3D0403023072311C301A06035504030C13437573746F6D2D4
   55220476C6F62616C204341310B3009060355040B0C02495431183016060355040A0
   C0F437573746F6D2D45522C20496E632E3111300F06035504070C0853616E204A6F7
   365310B300906035504080C024341310B3009060355040613025553301E170D32323
   13230363131333735395A170D3235313230353131333735395A30818D3131302F060
   35504030C28437573746F6D2D455220436F6D6D65726369616C204275696C64696E6
   7732052656769737472617231133011060355040B0C0A4F6666696365206F7073311
   83016060355040A0C0F437573746F6D2D45522C20496E632E310F300D06035504070
   C064F74746F7761310B300906035504080C024F4E310B30090603550406130243413
   059301306072A8648CE3D020106082A8648CE3D030107034200042030855BE846910
   B3ECB15CAF8572D3EB565BDC654A15EFAF6EDEBAA8B9E1608C5C910C93A20868FC85
   04D370F1B26FA9759F67983BB78863EDA9ACBEA5124F6A3818730818430090603551
   D1304023000300B0603551D0F0404030204F0301D0603551D0E04160414C9080B387
   D8DD85B3A59E7EC100B866393A9CA4C301F0603551D2304183016801492EA7640404
   A8FAB4F270BF3BC379D86CD7280F8302A0603551D250101FF0420301E06082B06010
   50507031C06082B0601050507030106082B06010505070302300A06082A8648CE3D0
   40302034900304602210091A2033692EB81503D53505FFC8DA326B1EE7DEA96F2917
   4F0B3341A07812201022100FF7339288108B712F418530A18025A895408CC45E0BB6
   78B46FBAB37DDB4D36B', h'30820243308201E9A00302010202082AEA0413A42DC1
   CE300A06082A8648CE3D0403023072311C301A06035504030C13437573746F6D2D45
   5220476C6F62616C204341310B3009060355040B0C02495431183016060355040A0C
   0F437573746F6D2D45522C20496E632E3111300F06035504070C0853616E204A6F73
   65310B300906035504080C024341310B3009060355040613025553301E170D323231
   3230363131333735395A170D3332313230333131333735395A3072311C301A060355
   04030C13437573746F6D2D455220476C6F62616C204341310B3009060355040B0C02
   495431183016060355040A0C0F437573746F6D2D45522C20496E632E3111300F0603
   5504070C0853616E204A6F7365310B300906035504080C024341310B300906035504
   06130255533059301306072A8648CE3D020106082A8648CE3D0301070342000497B1
   ED969164930985BBB8AC9A2AF9455CDFEEA4B11DE2E79D068BFA803926B4005251B3
   4F1C0815A4CBE03FBD1BBCB635F6431A22DE78653B87B99537ECE16CA3693067300F
   0603551D130101FF040530030101FF30250603551D11041E301C811A68656C704063
   7573746F6D2D65722E6578616D706C652E636F6D300E0603551D0F0101FF04040302
   0186301D0603551D0E0416041492EA7640404A8FAB4F270BF3BC379D86CD7280F830
   0A06082A8648CE3D0403020348003045022100D6D813B390BD3A7B4E85424BCB1ED9
   33AD1E981F2817B59083DD6EC1C5E3FADF02202CEE4406192BC767E98D7CFAE044C6
   807481AD8564A7D569DCA3D1CDF1E5E843']}, h'A11909C5A601020278183230323
   22D31322D30365432303A30343A31352E3735345A05581A041830168014CB8D98CA7
   4C51B58DDE7ACEF869A9443A8D666A6074823BFBBC9C2BCF2130958C9D28443A1012
   6A0587EA11909C5A40102074823BFBBC9C2BCF2130C585B3059301306072A8648CE3
   D020106082A8648CE3D030107034200042030855BE846910B3ECB15CAF8572D3EB56
   5BDC654A15EFAF6EDEBAA8B9E1608C5C910C93A20868FC8504D370F1B26FA9759F67
   983BB78863EDA9ACBEA5124F60D6D4A414441313233343536373839584068987DE8B
   007F4E9416610BBE2D48E1D7EA1032092B8BFCE611421950F45B22F17E214820C07E
   777ADF86175E25D3205568404C25FCEEC1B817C7861A6104B3D0D6D4A41444131323
   3343536373839', h'B1DD40B10787437588AEAC9036899191C16CCDBECA31C19785
   5CCB6BBA142D709FE329CBC3F76297D6063ACB6759EAB98E96EA4C4AA2135AA48A24
   7BAC1D6A3F'])



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C.5.  COSE-signed Voucher from MASA

   The resulting voucher is created by the MASA and returned to the
   Registrar:

     RES: 200 OK
       Content-Type: application/voucher+cose
       Body: <signed_voucher>

   The Registrar then returns the voucher to the Pledge:

     RES: 2.04 Changed
       Content-Format: 836 (application/voucher+cose)
       Payload: <signed_voucher>

   It is signed by the MASA's private key (see Appendix C.1.3) and can
   be verified by the Pledge using the MASA's public key that it stores.

   Below is the binary signed_voucher, encoded in hexadecimal (with lf
   added):

   D28443A10126A0590288A1190993A60102027818323032322D31322D30365432
   303A32333A33302E3730385A03F4074857EED786AD4049070859024730820243
   308201E9A00302010202082AEA0413A42DC1CE300A06082A8648CE3D04030230
   72311C301A06035504030C13437573746F6D2D455220476C6F62616C20434131
   0B3009060355040B0C02495431183016060355040A0C0F437573746F6D2D4552
   2C20496E632E3111300F06035504070C0853616E204A6F7365310B3009060355
   04080C024341310B3009060355040613025553301E170D323231323036313133
   3735395A170D3332313230333131333735395A3072311C301A06035504030C13
   437573746F6D2D455220476C6F62616C204341310B3009060355040B0C024954
   31183016060355040A0C0F437573746F6D2D45522C20496E632E3111300F0603
   5504070C0853616E204A6F7365310B300906035504080C024341310B30090603
   550406130255533059301306072A8648CE3D020106082A8648CE3D0301070342
   000497B1ED969164930985BBB8AC9A2AF9455CDFEEA4B11DE2E79D068BFA8039
   26B4005251B34F1C0815A4CBE03FBD1BBCB635F6431A22DE78653B87B99537EC
   E16CA3693067300F0603551D130101FF040530030101FF30250603551D11041E
   301C811A68656C7040637573746F6D2D65722E6578616D706C652E636F6D300E
   0603551D0F0101FF040403020186301D0603551D0E0416041492EA7640404A8F
   AB4F270BF3BC379D86CD7280F8300A06082A8648CE3D04030203480030450221
   00D6D813B390BD3A7B4E85424BCB1ED933AD1E981F2817B59083DD6EC1C5E3FA
   DF02202CEE4406192BC767E98D7CFAE044C6807481AD8564A7D569DCA3D1CDF1
   E5E8430B6D4A4144413132333435363738395840DF31B21A6AD3F5AC7F4C8B02
   6F551BD28FBCE62330D3E262AC170F6BFEDDBA5F2E8FBAA2CAACFED9E8614EAC
   5BF2450DADC53AC29DFA30E8787A1400B2E7C832

   The representation of signed_voucher in CBOR diagnostic format (with
   lf added) is:




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   18([h'A10126', {}, h'A1190993A60102027818323032322D31322D30365432303
   A32333A33302E3730385A03F4074857EED786AD4049070859024730820243308201E
   9A00302010202082AEA0413A42DC1CE300A06082A8648CE3D0403023072311C301A0
   6035504030C13437573746F6D2D455220476C6F62616C204341310B3009060355040
   B0C02495431183016060355040A0C0F437573746F6D2D45522C20496E632E3111300
   F06035504070C0853616E204A6F7365310B300906035504080C024341310B3009060
   355040613025553301E170D3232313230363131333735395A170D333231323033313
   1333735395A3072311C301A06035504030C13437573746F6D2D455220476C6F62616
   C204341310B3009060355040B0C02495431183016060355040A0C0F437573746F6D2
   D45522C20496E632E3111300F06035504070C0853616E204A6F7365310B300906035
   504080C024341310B30090603550406130255533059301306072A8648CE3D0201060
   82A8648CE3D0301070342000497B1ED969164930985BBB8AC9A2AF9455CDFEEA4B11
   DE2E79D068BFA803926B4005251B34F1C0815A4CBE03FBD1BBCB635F6431A22DE786
   53B87B99537ECE16CA3693067300F0603551D130101FF040530030101FF302506035
   51D11041E301C811A68656C7040637573746F6D2D65722E6578616D706C652E636F6
   D300E0603551D0F0101FF040403020186301D0603551D0E0416041492EA7640404A8
   FAB4F270BF3BC379D86CD7280F8300A06082A8648CE3D0403020348003045022100D
   6D813B390BD3A7B4E85424BCB1ED933AD1E981F2817B59083DD6EC1C5E3FADF02202
   CEE4406192BC767E98D7CFAE044C6807481AD8564A7D569DCA3D1CDF1E5E8430B6D4
   A414441313233343536373839', h'DF31B21A6AD3F5AC7F4C8B026F551BD28FBCE6
   2330D3E262AC170F6BFEDDBA5F2E8FBAA2CAACFED9E8614EAC5BF2450DADC53AC29D
   FA30E8787A1400B2E7C832'])

   In the above, the third element in the array is the voucher data
   encoded as a CBOR byte string.  When decoded, it can be represented
   by the following CBOR diagnostic notation:

   {2451: {1: 2, 2: "2022-12-06T20:23:30.708Z", 3: false, 7: h'57EED786
   AD404907', 8: h'30820243308201E9A00302010202082AEA0413A42DC1CE300A06
   082A8648CE3D0403023072311C301A06035504030C13437573746F6D2D455220476C
   6F62616C204341310B3009060355040B0C02495431183016060355040A0C0F437573
   746F6D2D45522C20496E632E3111300F06035504070C0853616E204A6F7365310B30
   0906035504080C024341310B3009060355040613025553301E170D32323132303631
   31333735395A170D3332313230333131333735395A3072311C301A06035504030C13
   437573746F6D2D455220476C6F62616C204341310B3009060355040B0C0249543118
   3016060355040A0C0F437573746F6D2D45522C20496E632E3111300F06035504070C
   0853616E204A6F7365310B300906035504080C024341310B30090603550406130255
   533059301306072A8648CE3D020106082A8648CE3D0301070342000497B1ED969164
   930985BBB8AC9A2AF9455CDFEEA4B11DE2E79D068BFA803926B4005251B34F1C0815
   A4CBE03FBD1BBCB635F6431A22DE78653B87B99537ECE16CA3693067300F0603551D
   130101FF040530030101FF30250603551D11041E301C811A68656C7040637573746F
   6D2D65722E6578616D706C652E636F6D300E0603551D0F0101FF040403020186301D
   0603551D0E0416041492EA7640404A8FAB4F270BF3BC379D86CD7280F8300A06082A
   8648CE3D0403020348003045022100D6D813B390BD3A7B4E85424BCB1ED933AD1E98
   1F2817B59083DD6EC1C5E3FADF02202CEE4406192BC767E98D7CFAE044C6807481AD
   8564A7D569DCA3D1CDF1E5E843', 11: "JADA123456789"}}





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   The largest element in the voucher is identified by key 8, which
   decodes to SID 2459 (pinned-domain-cert) based on the delta encoding
   defined by [RFC9254].  It contains the complete PKIX (DER-encoded
   X.509v3) certificate of the Registrar's domain CA.  This certificate
   is shown in Appendix C.2.3.

Appendix D.  Pledge Device Class Profiles

   cBRSKI allows implementers to select between various functional
   options for the Pledge, yielding different code size footprints and
   different requirements on Pledge hardware.  Thus for each product
   type an optimal trade-off between functionality, development/
   maintenance cost and hardware cost can be made.

   This appendix illustrates different selection outcomes by means of
   defining different example "profiles" of constrained Pledges.  In the
   following subsections, these profiles are defined and a comparison is
   provided.

D.1.  Minimal Pledge

   The Minimal Pledge profile (Min) aims to reduce code size and
   hardware cost to a minimum.  This comes with some severe functional
   restrictions, in particular:

   *  No support for EST re-enrollment: whenever this would be needed, a
      factory reset followed by a new onboarding process is required.

   *  No support for change of Registrar: for this case, a factory reset
      followed by a new onboarding process is required.

   This profile would be appropriate for single-use devices which must
   be replaced rather than re-deployed.  That might include medical
   devices, but also sensors used during construction, such as concrete
   temperature sensors.

D.2.  Typical Pledge

   The Typical Pledge profile (Typ) aims to support a typical cBRSKI
   feature set including EST re-enrollment support and Registrar
   changes.










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D.3.  Full-featured Pledge

   The Full-featured Pledge profile (Full) illustrates a Pledge category
   that supports multiple onboarding methods, hardware real-time clock,
   BRSKI/EST resource discovery, and CSR Attributes request/response.
   It also supports most of the optional features defined in this
   specification.

D.4.  Comparison Chart of Pledge Classes

   The below table specifies the functions implemented in the three
   example Pledge classes Min (Appendix D.1), Typ (Appendix D.2) and
   Full (Appendix D.3).






































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    +============================================+=======+=====+======+
    | Functions Implemented                      |  Min  | Typ | Full |
    +============================================+=======+=====+======+
    | *General*                                  |       |     |      |
    +--------------------------------------------+-------+-----+------+
    | Support cBRSKI onboarding                  |   Y   |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Support other onboarding method(s)         |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Real-time clock and cert time checks       |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | *cBRSKI*                                   |       |     |      |
    +--------------------------------------------+-------+-----+------+
    | CoAP discovery for rt=brski*               |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Support pinned Registrar public key (RPK)  |   Y   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Support pinned Registrar certificate       |   -   |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Support pinned Domain CA                   |   -   |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | *EST-coaps*                                |       |     |      |
    +--------------------------------------------+-------+-----+------+
    | Explicit TA database size (#certs)         |   0   |  3  |  8   |
    +--------------------------------------------+-------+-----+------+
    | CoAP discovery for rt=ace.est*             |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | GET /att and response parsing              |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | GET /crts format 62 (multiple CA certs)    |   -   |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | GET /crts format 281 (multiple CA certs)   |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | ETag handling support for GET /crts        |   -   |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Re-enrollment supported                    | - (*) |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Section 6.5.1 optimized procedure          |   Y   |  Y  |  -   |
    +--------------------------------------------+-------+-----+------+
    | Pro-active re-enrollment at own initiative |   -   |  -  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Periodic trust anchor retrieval GET /crts  | - (*) |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+
    | Supports change of Registrar identity      | - (*) |  Y  |  Y   |
    +--------------------------------------------+-------+-----+------+

       Table 5: Comparison Chart of Pledge Classes Min, Typ and Full




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   Notes: (*) means only possible via a factory-reset followed by a new
   cBRSKI onboarding procedure.

Appendix E.  Pledge Discovery of Onboarding and Enrollment Options

   The discovery functionality described in this section is informative
   only: it derives from the normative CoRE documents [RFC6690],
   [RFC7252] and from [RFC9148].  In typical cases, for a constrained
   Pledge that only supports a single onboarding and enrollment method,
   this functionality is not needed.

   Note that the full-featured Pledge class defined in Appendix D.3 does
   support CoAP discovery functionality.

E.1.  Pledge Discovery Query for All cBRSKI Resources

   A Pledge that wishes to discover the available cBRSKI onboarding
   options/formats can do a discovery operation using CoAP discovery per
   Section 7 of [RFC7252] and Section 4 of [RFC6690].  It first sends a
   CoAP discovery query to the Registrar over the secured DTLS
   connection.  The Registrar then responds with a CoRE Link Format
   payload containing the requested resources, if any.

   For example, if the Registrar supports a cBRSKI base resource /b in
   addition to the longer /.well-known/brski base resource, and supports
   only the voucher format application/voucher+cose (836), and status
   reporting in both CBOR (60) and JSON (50) formats, a CoAP resource
   discovery request and response may look as follows:

     REQ: GET /.well-known/core?rt=brski*

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         </b>;rt=brski,
         </b/rv>;rt=brski.rv;ct=836,
         </b/vs>;rt=brski.vs;ct="50 60",
         </b/es>;rt=brski.es;ct="50 60"

   In this case, the Registrar returns only the shorter URI paths
   matching the query filter, which are located under the /b base
   resource.  The /.well-known/brski based URI paths are not returned,
   as these are assumed to be well-known (i.e. mandatory to support for
   a Registrar that offers this functionality under /b.)







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   The Registrar is however under no obligation to provide the shorter
   URLs under /b, and may respond to this query with only the /.well-
   known/brski/\<short-name\> resources for the short names as defined
   in Table 1, if these resources are not hosted anywhere else.  This
   case is shown in the below interaction:

     REQ: GET /.well-known/core?rt=brski*

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         </.well-known/brski>;rt=brski,
         </.well-known/brski/rv>;rt=brski.rv;ct=836,
         </.well-known/brski/vs>;rt=brski.vs;ct="50 60",
         </.well-known/brski/es>;rt=brski.es;ct="50 60"

   When responding to a discovery request for cBRSKI resources, the
   Registrar may return the full resource paths for all <short-name>
   resources and the content-formats supported by these resources (using
   ct attributes) as shown in the above examples.  This is useful when
   multiple content-formats are supported for a particular resource on
   the Registrar and the discovering Pledge also supports multiple.

   Registrars that have implemented any cBRSKI or EST-coaps URI paths
   outside of /.well-known must process a request on the corresponding
   /.well-known/brski/\<short-name\> or /.well-known/est/\<short-name\>
   URI paths identically.  In particular, a Pledge may use the longer
   (well-known) and shorter URI paths in any combination.

   A Registrar may also be implemented without support for the
   (optional) CoAP discovery.  In that case, it may for example return a
   4.04 Not Found as shown in the example below, in case the Registrar
   does not host the resource /.well-known/core at all.  In such case,
   the Pledge cannot discover any onboarding/enrollment options and so
   it has to rely on the default cBRSKI resources under /.well-known/
   brski/... and /.well-known/est/....

     REQ: GET /.well-known/core?rt=brski*

     RES: 4.04 Not Found











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E.2.  Pledge Discovery Query for the cBRSKI Base Resource

   In case the client queries for only rt=brski type resources, the
   Registrar responds with only the base path for the cBRSKI resources
   (rt=brski, resource /b in earlier examples) and no others.  (So, the
   query is "rt=brski", without the wildcard character.)  This is shown
   in the below example.  The Pledge in this case requests only the
   cBRSKI base resource of type rt=brski to check if cBRSKI is supported
   by the Registrar and if a shorter-length cBRSKI base resource path is
   supported or not.  In this case, the Pledge is not interested to
   check what voucher request formats, or status telemetry formats --
   other than the mandatory default formats -- are supported.  The
   compact response below then shows that the Registrar indeed supports
   a cBRSKI resource at /b:

     REQ: GET /.well-known/core?rt=brski

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         </b>;rt=brski

   The Pledge can now start using any of the cBRSKI resources
   /b/\<short-name\> in a next CoAP request to the Registrar.  In above
   example, again the well-known resource present under /.well-known/
   brski is not returned because this is assumed to be well-known to the
   Pledge and mandatory to support for a Registrar that offers this
   functionality under /b.

   As a follow-up example, the Pledge can now start the onboarding by
   sending its PVR:

     REQ: POST /b/rv
       Content-Format: 836 (application/voucher+cose)
       Accept: 836 (application/voucher+cose)
       Payload: <binary COSE-signed PVR>

E.3.  Usage of ct Attribute

   The return of multiple content-formats in the 'ct' link format
   attribute by the Registrar allows the Pledge to choose the most
   appropriate one for a particular operation, and allows extension with
   new voucher formats.  Note that only content-format 836 (application/
   voucher+cose) is defined in this document for the voucher request
   resource (/rv), both as request payload and as response payload.  If
   the 'ct' attribute is not indicated for the /rv resource in the CoRE
   Link Format description, this implies that at least format 836 is
   supported and maybe more.



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   Note that this specification allows for application/voucher+cose
   payloads to be transmitted over HTTPS, as well as for application/
   voucher-cms+json and other formats yet to be defined over CoAP.  The
   burden for this flexibility is placed upon the Registrar.  A Pledge
   on constrained hardware is expected to support a single format only.

   The Pledge needs to support one or more formats for the PVR and
   resulting voucher.  The MASA needs to support all formats that the
   associated Pledges use.

   In the below example, a Pledge queries specifically for the brski.rv
   resource type to learn what voucher formats are supported:

     REQ: GET /.well-known/core?rt=brski.rv

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         </b/rv>;rt=brski.rv;ct="836 65123 65124"

   The Registrar returns 3 supported voucher formats: 836, 65123, and
   65124.  The first is the mandatory application/voucher+cose.  The
   other two are numbers from the Experimental Use number range of the
   CoAP Content-Formats sub-registry, which are used as mere examples.
   The Pledge can now make a selection between the supported formats.

   Note that if the Registrar only supports the default content-formats
   for each cBRSKI resource as specified by this document, it may omit
   the ct attributes in the discovery query response.  For example as in
   the following interaction:

     REQ: GET /.well-known/core?rt=brski*

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         </b>;rt=brski,
         </b/rv>;rt=brski.rv,
         </b/vs>;rt=brski.vs,
         </b/es>;rt=brski.es

E.4.  EST-coaps Resource Discovery

   The Pledge can also use CoAP discovery to identify enrollment
   options, for example enrollment using EST-coaps or other methods.
   The below example shows a Pledge that wants to identify EST-coaps
   enrollment options by sending a discovery query.  This is done either
   before or after the voucher has been validated.



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     REQ: GET /.well-known/core?rt=ace.est*

     RES: 2.05 Content
       Content-Format: 40 (application/link-format)
       Payload:
         </e/crts>;rt=ace.est.crts;ct="62 281 287",
         </e/sen>;rt=ace.est.sen;ct="281 287",
         </e/sren>;rt=ace.est.sren;ct="281 287",
         </e/att>;rt=ace.est.att,
         </e/skg>;rt=ace.est.skg,
         </e/skc>;rt=ace.est.skc

   The response from the Registrar indicates that EST-coaps enrollment
   (/sen) and re-enrollment (/sren) is supported, with a choice of two
   content-formats for the response payload: either a PKCS#7 container
   with a single LDevID certificate (application/pkcs7-mime;smime-
   type=certs-only, content-format 281) which is the BRSKI [RFC8995]
   encoding, or just a single LDevID certificate (application/pkix-cert,
   content-format 287) which is the default cBRSKI encoding.

   For the EST cacerts resource (/crts) there are three content-formats
   supported: an application/multipart-core container (62) per
   Section 6.5.5, a PKCS#7 container with all CA certificates (287), or
   a single (most relevant) CA certificate (281).

   The Pledge can now send a CoAP request to one of the discovered
   resources, with the Accept Option to indicate which return payload
   content-format the Pledge wants to receive.

Acknowledgements

   We are very grateful to Jim Schaad for explaining COSE/CMS choices
   and for correcting early versions of the COSE_Sign1 objects.

   Michel Veillette did extensive work on _pyang_ to extend it to
   support the SID allocation process, and this document was among its
   first users.

   Russ Housley , Daniel Franke , Henk Birkholtz , Kathleen Moriarty ,
   Xufeng Liu and Karl Moberg provided review feedback.

   The BRSKI design team has met on many Tuesdays and Thursdays for
   document review.  The team includes the authors and: Aurelio
   Schellenbaum , David von Oheimb , Steffen Fries , Thomas Werner ,
   Bill Atwood and Toerless Eckert .

   Darrel Miller , Orie Steele and Manu Sporny provided review feedback
   on the registration of the +cose structured syntax suffix.



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   Carsten Bormann suggested the use of CBOR Web Token (CWT) claims in
   the voucher's COSE header.

Changelog

   -30: Require Pledge's DTLS cert chain to be included in RVR 'x5bag'
   (#343).  Add support for CoAP Uri-Path-Abbrev Option (#336).  Move
   'idevid-issuer' clarification text to draft-8366bis.  Update the
   duplicate serial number attack to focus only on the case where the
   attack could be successful (equal CAs).  Update section references
   draft-ietf-anima-8366bis to latest version.  Remove reference to the
   to-be-deprecated RFC 8366.  Align terms and notation with draft-ietf-
   anima-8366bis.  Editorial (wording) updates.

   -29: Clarify that each brski.jp link indicates a root resource (/)
   (#335).  Clarify that Pledge uses IP link-local address of JP's
   discovery response, instead of the UTF-8 encoded IP address literal
   (#334).  Add example of Join Proxy offering multiple Registrars
   (endpoints) (#333).  Updated CoAP request/response formatting of
   examples.  Updated acknowledgements (#331).  Editorial updates.

   -28: Cleanup of normative/informative references, setting each to
   right category.  Bugfix and clarification in text around EdDSA Curve
   selection.  Added section on additional information in COSE header
   with 'iat' CWT timestamp example.  Updates to BRSKI Well-Known URIs
   registry, including a rename of the "URI" column (#326).  Unify COSE
   header parameters terminology (#330).  Text formatting and editorial
   updates.

   -27: Clarify x5bag for storing signing chain and Registrar removes
   unprotected x5bag/x5chain (#324, #323, #230).  Clarify RPK use with
   "placeholder" certificate.  Merged the very similar BRSKI-MASA
   security considerations sections (#312).  Require CBOR format for
   Pledge's/EST-client's telemetry (#309, #317).  Removed figure
   captions from code examples for consistency (#315).  Add base
   resource type (rt) for "ace.est" and related terminology (#314).
   Update IEEE 802.1AR reference to 2018 version (#313).  Editorial
   updates.

   -26: Updated I-D/RFC references to newer versions.  Corrected "sub-
   registry" term to official "registry", in IANA section.  Explicitly
   imported terminology from [RFC7252].  Corrected "router" term in
   Thread/MLE section, with clarifications, and [Thread] reference fix.
   Moved references between "Informative" and "Normative" based on
   what's required to implement all the optional features.  Removal of
   some lingering legacy text.  Editorial improvements, bugfixes and
   typo corrections.




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   -25: Moved all software/library support info into Appendix A and
   added "open source" section; Removed use of formal Extends/Amends
   Update-tags (#303, #304); Moved Section 14 to Appendix E (#302);
   Editorial improvements.

   -24: Rephrased well-known URL requirement in 14.1 (#292, #293); Added
   paragraph on future certificate formats like C509 (#281, #294); Add
   formal specification for CoAP discovery of Join Proxy by Pledge,
   instead of only showing examples (#296, #300); Enable mDNS discovery
   of Join Proxy by Pledge (also in mesh networks) and list service name
   to use (#297, #299); Add requirement to support content-format 287 in
   /sen and /sren response (#295, #298).

   -23: Removed Update tag for RFC 8366 (#285, #288); Introduced cBRSKI
   acronym (#284, #286); Added Update tag for RFC 9148 (#283, #289);
   Keep CoAP discovery as only mechanism and refer to future discovery
   work (#279, #282, #290); Introduce formal CBOR diagnostics ellipsis
   elision syntax (#281, #287); Support for multi-tier CAs by
   introducing multipart-core /crts format (#275, #291); Terminology
   updated for consistency with RFC 8366-bis (#274, #280); Rename
   voucher media type to application/voucher+cose and register +cose SSS
   (#264, #277); Editorial changes including section restructuring.

   -22: Streamlined text to focus mostly on the default flow, with
   optional functions moved to their own sections (#269, #273); For DTLS
   1.3 client, use the record_size_limit extensions RFC 8449 (#270);
   Editorial updates; Reference rfc6125bis updated to RFC 9525.

   -11 to -21: (For change details see GitHub issues https://github.com/
   anima-wg/constrained-voucher/issues , related Pull Requests and
   commits.)

   -10: Design considerations extended; Examples made consistent.

   -08: Examples for cose_sign1 are completed and improved.

   -06: New SID values assigned; regenerated examples.

   -04: voucher and request-voucher MUST be signed; examples for signed
   request are added in appendix; IANA SID registration is updated; SID
   values in examples are aligned; signed cms examples aligned with new
   SIDs.

   -03: Examples are inverted.







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   -02: Example of requestvoucher with unsigned application/cbor is
   added; attributes of voucher "refined" to optional; CBOR
   serialization of vouchers improved; Discovery port numbers are
   specified.

   -01: application/json is optional, application/cbor is compulsory;
   Cms and cose mediatypes are introduced.

   -00: Initial version.

Authors' Addresses

   Michael Richardson
   Sandelman Software Works
   Email: mcr+ietf@sandelman.ca


   Peter van der Stok
   vanderstok consultancy
   Email: stokcons@kpnmail.nl


   Panos Kampanakis
   Cisco Systems
   Email: pkampana@cisco.com


   Esko Dijk
   IoTconsultancy.nl
   Email: esko.dijk@iotconsultancy.nl





















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