- Basic goal of 1-of-\(N\) ring signatures
- Recap: the \(\Sigma\)-protocol
- OR of \(\Sigma\)-protocols, CDS 1994
- Abe-Ohkubo-Suzuki (AOS) 2002 (broken version)
- Security weaknesses
- Key prefixing
- Borromean, Maxwell-Poelstra 2015
- Linkability and exculpability
- AND of \(\Sigma\)-protocols, DLEQ
- Liu-Wei-Wong 2004
- Security arguments for the LWW LSAG
- Back 2015; compression, single-use
- Fujisaki-Suzuki 2007 and Cryptonote 2014
- Monero MLSAG

The idea of a ring signature (the term itself is a bit sloppy in context, but let's stick with it for now) is simple enough:

An owner of a particular private key \(x\) signs a message \(m\) by taking, usually without setup or interaction, a whole set of public keys, one of which is his (\(P=xG\)), and forms a signature (exact form unspecified) such that there is proof that **at least one** of the private keys is known to the signer, but which one was responsible for the signature is not known by the verifier, and not calculatable.

Obviously that's pretty vague but captures the central idea. We often use the term "ring" because the construction must have some symmetry over the entire set of \(n\) public keys, and a ring/circle represents symmetry of an arbitrarily high order (limit of an \(n\)-gon). Less abstractly it could be a good name because of some "loop"-ing aspect of the algorithm that constructs the signature, as we'll see.

What properties do we want then, in summation?

- Unforgeability
- Signer ambiguity

We may want additional properties for some ring signatures, as we'll see.

In the following sections I want to cover some of the key conceptual steps to the kinds of ring signatures currently used in cryptocurrency protocols; most notably Monero, but also several others; and also in the Confidential Transactions construction (see: Borromean ring signatures, briefly discussed here). I will also discuss security of such constructions, in much less detail than the previous blog (on the security of Schnorr signatures), but showing how there are several tricky issues to be dealt with, here.

We consider a prover \(\mathbb{P}\) and a verifier \(\mathbb{V}\).

A \(\Sigma\)-protocol is a three step game, in which the prover convinces the verifier of something (it can be \(\mathbb{P}\)'s knowledge of a secret, but it can also be something more complicated), in zero knowledge. Readers interested in a much more detailed discussion of the logic behind this and several applications of the idea can read Sections 3 and 4 of my From Zero (Knowledge) to Bulletproofs writeup, especially section 4.1.

In brief, the three step game is:

\(\mathbb{P} \rightarrow \mathbb{V}\): **commitment**

\(\mathbb{V} \rightarrow \mathbb{P}\): **challenge**

\(\mathbb{P} \rightarrow \mathbb{V}\): **response **

A few minor notes on this: obviously the game is not literally over with the response step; the verifier will examine the response to establish whether it is valid or invalid.

The **commitment** will usually in this document be written \(R\) and will here always be a point on an elliptic curve, which the prover may (or may not! in these protocols) know the corresponding scalar multiple (private key or nonce) \(k\) such that \(R=kG\).

The **challenge** will usually be written \(e\) and will usually be formed as the hash of some transcript of data; the subtleties around exactly *what* is hashed can be vitally important, as we'll see. (This is in the "Fiat-Shamir transform" case; we discussed the pure interactive challenge case a bit in the previous blog and many other places!)

The **response** will usually be a single scalar which will usually be denoted \(s\).

We will be playing with this structure a lot: forging transcripts \(R, e, s\); running multiple instances of a \(\Sigma\)-protocol in parallel and performing logical operations on them. All of this will play out *mostly* in the form of a Schnorr signature; again, refer to previous blog posts or elementary explanations (including those written by me) for more on that.

Let's start with the OR of \(\Sigma\)-protocols. I *believe* this solution is due to Cramer, Damgård and Schoenmakers '94

(Historical note: the "believe" is because I've seen it cited to that paper (which is famous for good reason, I guess); but in the paper they actually attribute *this specific idea* to "M. Ito, A. Saito, and T. Nishizeki: Secret Sharing Scheme realizing any Access Structure, Proc. Glob.Com. (1987)" ; unfortunately I can't find that on the 'net).

It is also described, with a brief discussion of its security proof, in Boneh-Shoup Sec 19.7.2.

This is not, as far as I know, used at all(?) nor that widely discussed, but it is in some sense the most simple and logical way to get a 1 out of \(N\) ring signature; use the XOR (\(\oplus\)) operation:

We have in advance a set of public keys \(P_i\). We only know one private key for index \(j\), \(x_j\).

We'll now use a standard three move \(\Sigma\)-protocol to prove knowledge of **at least one key** without revealing which index is \(j\).

We're going to fake the non-\(j\)-index signatures in advance. Choose \(s_i \stackrel{$}{\leftarrow} \mathbb{Z}_N\ ,\ e_i \stackrel{$}{\leftarrow} \mathbb{Z}_N \quad \forall i \neq j\).

Calculate \(R_i = s_iG - e_iP_i \quad \forall i \neq j\).

For the real signing index, \(k_j \stackrel{$}{\leftarrow} \mathbb{Z}_N\ ,\quad R_j = k_jG\).

We now have the full set of commitments: \((R_i \ \forall i)\)

Now for the clever part. In an interactive \(\Sigma\)-protocol, we would at this point receive a random challenge \(e \in \mathbb{Z}_N\). For the Fiat Shamir transformed case, noninteractively (as for a signature), we use the constructed \(R\)-values as input to a hash function, i.e. \(e = H(m||R_i \ldots)\). We have already set the non-signing index \(e\)-values, for the signing index we set \(e_j = e \oplus (\bigoplus_{i \ne j}{e_i})\).

This allows us to calculate \(s_j = k_j + e_j x_j\), and we now have the full set of 'responses' for all the \(\Sigma\)-protocols: \(s_i \ \forall i\). (but here we are using Fiat Shamir, so it's not actually a response).

By working this way we have ensured that the signature verifier can verify that the logical XOR of the three \(e\)-values is equal to the Fiat Shamir based hash-challenge, e.g. for the case of three "signatures", we will have:

\(e = e_1 \oplus e_2 \oplus e_3 \stackrel{?}{=} H(m||R_0||R_1||...)\)

where the verifier would calculate each \(R_i\) as \(s_iG - e_iP_i\).

The excellent feature of this of course is that it is perfectly hidden which of the three indexes was genuine. But the bad news is that the protocol as stated, used let's say as a signature scheme, requires about twice as many field elements as members of the group of signers. The verifier needs to be given \((s_1, \ldots s_n),(e_1 \ldots e_n)\).

Another excellent feature: this is not restricted to the Schnorr ID protocol. It can work with another identity protocol, and even better, it could work with a *mix* of them; they only have to share the one challenge \(e\).

This is an excellent paper generally, but its stand-out contribution, in this context, is a **more compact** version of the 1 of n ring signature above. To clarify here, both this and the previous are \(O(n)\) where \(n\) is the group size, so "much more compact" is about the constant factor (scale not scaling!); we reduce it from roughly 2 to roughly 1.

"Broken version" - here I'll present a slightly simpler form than the one in the paper, and then explain the serious problem with it - which I hope will be productive. **Please don't mistake this as meaning that the AOS design was broken, it was never presented like this in the paper!**

Anyway, I think the best explanation for what's going on here conceptually is due to A. Poelstra in the Borromean ring signatures paper, particularly Section 2 ; the reference to time travel may seem whimsical but it gets to the heart of what's going on here; it's about having a simulated form of causality with one way functions, and then violating that.

In short: creating an ordinary Schnorr sig without the key (i.e. forging) is impossible because, working at the curve point level of the equation (\(sG = R + H(m||R)P\)), you need to know the hash value before you can calculate \(R\), but you need to know the value of \(R\) before you can calculate the hash. So we see that two one way functions are designed to conflict with one another; only by removing one of them (going from curve points to scalar eqn: (\(s = k + H(m||kG)x\)), can we now create a valid \(s, R, m\) set.

To achieve that goal over a set of keys, we can make that "simulated causality enforcement" be based on the same principle, but over a set of equations instead of one. The idea is to make the commitment \(H(m||R)\) use the \(R\) value from the "previous" signer/key/equation, where "previous" is modulo \(N\), i.e. there is a loop of dependencies (a ring, in fact).

Quick description:

Our goal is a list of \(N\) correctly verifying Schnorr signature equations, with the tweak as mentioned that each hash-value refers to the "previous" commitment. We will work with \(N=4\) and index from zero for concreteness. Our goal is:

\(s_0 G = R_0 + H(m||R_3)P_0\)

\(s_1 G = R_1 + H(m||R_0)P_1\)

\(s_2 G = R_2 + H(m||R_1)P_2\)

\(s_3 G = R_3 + H(m||R_2)P_3\)

Again for concreteness, we imagine knowing specifically the private key \(x_2\) for index 2, only. We can successfully construct the above, but only in a certain sequence:

Choose \(k_2 \stackrel{$}{\leftarrow} \mathbb{Z}_N,\ R_2 = k_2G\), choose \(s_3 \stackrel{$}{\leftarrow} \mathbb{Z}_N\).

\(\Rightarrow R_3 = s_3 G - H(m||R_2)P_3\). Now choose \(s_0 \stackrel{$}{\leftarrow} \mathbb{Z}_N\).

\(\Rightarrow R_0 = s_0 G - H(m||R_3)P_0\). Now choose \(s_1 \stackrel{$}{\leftarrow} \mathbb{Z}_N\).

\(\Rightarrow R_1 = s_1 G - H(m||R_0)P_1\).

Last, do not choose but **calculate** \(s_2\): it must be \(s_2 = k_2 + H(m||R_1)x_2\).

After this set of steps, the set of data: \(e_0, s_0, s_1, s_2, s_3\) can be verified without exposing which private key was known. Here is the verification:

Given \(e_0, s_0\), reconstruct \(R_0 = s_0G -e_0P_0\).

\(\Rightarrow e_1 =H(m||R_0)\ ,\ R_1 = s_1 G - e_1P_1\)

\(\Rightarrow e_2 =H(m||R_1)\ ,\ R_2 = s_2 G - e_2P_2\)

\(\Rightarrow e_3 =H(m||R_2)\ ,\ R_3 = s_3 G - e_3P_3\)

**Check**: \(e_0 \stackrel{?}{=} H(m||R_3)\).

The description above can't be described as secure.

To give a hint as to what I mean: is there something **not completely fixed** in the above construction? Maybe an issue that's not even specific to the "ring" construction, but even for any one of the signature equations?

....

The answer is the keys, \(P_i\). We can in the most general case consider three scenarios, although there may be some gray areas between them:

- Key(s) fixed in advance: \(P_1 \ldots P_N\) are all specified before doing anything, and not allowed to change by the verifier. Every signature must be on that set of keys.
- The
*set**of possible keys*is fixed in advance exactly as described above, but the*set of keys used in the ring*is chosen by the signer, dynamically, in signing oracle queries or forgery attempts. - Even the set of possible keys is dynamic. That is to say, any valid curve point (for EC case) is a valid potential key in (ring) signature.

This is not a full taxonomy of possible attack scenarios, either. Not only must we consider the difference between EUF-CMA and SUF-CMA as was discussed in the previous blog (a reminder: with SUF, a forger should not be able to even create a second signature on the same message - ECDSA doesn't have this in naive form), but much more: we must also consider which of the above three key settings applies.

Even outside of ring signature settings, just considering a large scale deployment of a signature scheme across millions or billions of keys, could mean that the difference between these cases really matters. In this paper by Dan Bernstein the term MU-UF-CMA is used to refer to the "multi-user" setting for this, where only single-key signatures are used but one must consider whether having billions of other keys and signing oracles for them might impact the security of **any one** key (notice the huge difference between "I want to forge on \(P\)" and "I want to forge on any existing key" is, in this scenario).

So enough about settings, what exactly constitutes a security problem with the above version of the AOS ring sig?

Consider any one element in the ring like:

\(s_0 = R_0 + H(m||R_3)P_0\)

where, for concreteness, I choose \(n=4\) and look at the first of 4 signature equations. Because of Schnorr's linearity (see this earlier blog post for some elucidations on the *advantage* of this linearity, although it was also noted there that it had concomitant dangers (worse, actually!)), there are two obvious ways we could tweak this equation:

(1) Tweaked \(s\) values on fixed message and tweaked keys:

Choose \(\alpha \in \mathbb{Z}_N\) and set \(s' = s_0 +\alpha\). We will not alter \(R=kG\), but we alter \(P_0 \rightarrow P_0 + e_0^{-1}\alpha G\). This makes the verification still work **without altering the fixing of the nonce in the hash value \(e_0\):**

\(s_0 G + \alpha G = R_0 + e_0 P_0 + \alpha G = R_0 + e_0\left(P_0 + e_0^{-1}\alpha G\right)\)

So it's really not clear how bad this failing is; it's *kinda* a failure of strong unforgeability, but that notion doesn't precisely capture it: we created a new, valid signature against a new key, but with two severe limitations: we weren't able to alter the message, and also, we weren't able to *choose* the new key \(P'\). That last is slightly unobvious, but crucial : if I have a pre-prepared \(P^{*}\), I cannot choose \(\alpha\) to get \(P' = P^{*}\) as that would require a discrete logarithm break.

A final statement, hopefully obvious: the above can apply to any and all of the elements of the ring, so the forgery could consist of an entirely different and random set of keys, not related to the starting set; but the message would be the same, as would the \(R\) values.

(2) Completely different messages on tweaked keys, with the same signature

This one is almost certainly more important. Algebraically, we here allow alterations to the \(e\) values, using multiplication rather than addition:

Given the same starting \(s_0\) as in (1), we take a chosen new message \(m^{*}\) and calculate the new \(e^{*} = H(m^{*}||R_3)\). If we likewise tweak the public key we get that \(s_0, R_0\) is a valid signature on the new message, with the tweaked key:

\(s_0 G = R_0 + e_0^{*}\left(\frac{e_0}{e_0^{*}} P_0\right)\)

We can see here that this produces a forgery with the same signature values (but different hash values) on the new keys.

Most definitions of security against forgery require the attacker to create a signature on a not-previously-queried message - so this *is* a successful attack, by most measures.

However it does share the same limitation with (1) mentioned above - that you cannot "control" the keys on which you get a signature, unless you know a relative discrete log between one of the existing keys and your new key, which implies you knew the secret key of the first (in which case all this is pointless; whenever you have a private key, there is no forgery on it).

**All of this should make very clear the reason why the real AOS (see Section 5.1 of the paper) discrete-log ring signature fixes the entire set of keys inside the hash, i.e. \(e_i = H(m || R_{(i-1)\%n}|| P_0 \ldots P_{n-1})\).**

The method in the previous bolded sentence is sometimes called "key-prefixing". One way of looking at it: the Fiat-Shamir transform that takes the Identity Protocol into a signature scheme, should hash the conversation transcript between the prover and verifier, previous to the challenge step; by including the public keys in this hash, we are treating the keyset as part of the conversation transcript, rather than something ex-protocol-run.

Also, the discussion above (both cases (1) and (2)) show clearly that the same weakness exists for a single (\(n=1\)) key case.

And yet, for the single key case, it was not a done deal historically - this caused real world arguments!. After all, there are many use cases where the key *is* a given ex-protocol-run, plus there may be some practical disadvantage to doing the key-prefixing.

In this paper from CRYPTO-2016, the controversy arising out of this is elucidated, showing that these theoretical concerns had very substantial impact on arguably the largest real world crypto usage (TLS):

"Key-prefixing comes with the disadvantage that the entire public-key has to

be available at the time of signing. Specifically, in a CFRG message from Sep-

tember 2015 Hamburg [32] argues “having to hold the public key along with

the private key can be annoying” and “can matter for constrained devices”.

Independent of efficiency, we believe that a cryptographic protocol should be

as light as possible and prefixing (just as any other component) should only

be included if its presence is justified. Naturally, in light of the GMLS proof,

Hamburg [32] and Struik [44] (among others) recommended against key prefixing

for Schnorr. Shortly after, Bernstein [10] identifies the error in the GMLS theo-

rem and posts a tight security proof for the key-prefixed variant of Schnorr signa-

tures. In what happens next, the participant of the CFRG mailing list switched

their minds and mutually agree that key-prefixing should be preferred, despite of

its previously discussed disadvantages. Specifically, Brown writes about Schnorr

signatures that “this justifies a MUST for inclusion of the public key in the mes-

sage of the classic signature” [16]. As a consequence, key-prefixing is contained in

the current draft for EdDSA [33]..."

*Technical note: the "GMLS proof" mentioned in the above is the proof given in this paper, that was intended to reduce the security of the multi-user setting to that of the single-user setting, and that Dan Bernstein's paper previously mentioned proved to be invalid.*

What's the TLDR? Fix the keys in any group/ring/multisignature. And even that may not be enough, see MuSig for details of why it really isn't, in the scenario of Bitcoin aggregated multisig.

I covered this extensively (including description of AOS as above) in my CT writeup section 3.2

The idea of the construction as outlined in the paper by Maxwell, Poelstra is to increase the space-efficiency of the published proof even more. By having several ring signatures joined at a single index we get a reduction in the number of \(e\) values we publish. This is basically the same idea as the "AND of \(\Sigma\)-protocols" discussed a little later in this document (although here we will only be using it for achieving a specific goal, "Linkability", see more on this next).

For the real world context - Borromean ring signatures are used in certain implementations of Confidential Transactions (e.g. Liquid by Blockstream) today, and were previously used also in Monero for the same goal of CT. They are a radically different use-case of ring signatures to the one mostly described in the below; instead of using a ring signature to hide the identity of a signer, they are used to hide which exponent contains values in the encoding of a value committed to in a Pedersen commitment. This allows arithmetic to be done on the Pedersen-committed amount without worrying about overflow into negative values modulo \(N\).

In this section we'll briefly describe certain key features that turn out to be useful in some real-world applications of a ring signature, before in the following sections laying out how these features are, or are not, achieved.

At first glance, the idea "linkability" with a ring signature seems to be a contradiction. Since we are trying to achieve signer ambiguity/anonymity, we don't really want any "linking" being done. But the idea is rather clever, and proves to be very interesting for digital cash.

In a **linkable** ring signature, a participant with key \(P \in L\) (i.e. \(L\) is a particular set of public keys), should be able to produce one ring signature on a given message, but should not be able to do so again without the two ring signatures being linked. Thus, functionally, each participant can only make such a signature once (note: they can still retain anonymity if double-signing).

This restriction-to-one-signature-while-keeping-anonymity is easily seen to be valuable in cases like electronic voting or digital cash, as well as the oft-cited example explained in the next paragraph.

The **spontaneity** property should be a lot more obvious. Consider the example of a whistleblower. We would want individuals in some large group (e.g. government bureaucrats) to attest to a statement, while only revealing group membership and not individual identity. Clearly this is not workable if it requires cooperation of other members of the group (even in any setup phase), so it's necessary that the individual can create the ring signature "spontaneously", knowing only the public key of other participants.

The paper uses the abbreviation LSAG for this type of signature: "Linkable Spontaneous Anonymous Group" signature.

Note that the previous two constructions (CDS, AOS) can also have this spontaneity property; but not the linkability property.

A ring signature can be described as exculpable if, even given knowledge of the signing private key, an adversary cannot deduce that that signing key was the one used to create the ring signature.

Notice that such a property may be immensely important in a range of scenarios where a ring sig is useful - e.g. for a whistleblower whose cryptographic keys were stolen or extracted by force, he could still plausibly deny being the origin of a leak.

The reader can easily verify that the AOS construction, for example, has this exculpability. The fact that a particular key is released e.g. \(x_2\) in our concrete example, does not allow inference of it having been used to create that signature. Any other key could have created the signature, using the same signing algorithm.

The LWW LSAG, which we'll describe shortly, is on the other hand **culpable**, i.e. the opposite - because the key image can be verified to be tied to one particular key.

It's easy to see that the two properties **exculpability** and **linkability** are somewhat in conflict, although I'm not aware of a theorem that *absolutely requires* linkability to somehow tag one key in case it is leaked.

Lastly, I'll mention **claimability**, which is briefly described also in the LWW paper (see below). It may be possible for the owner of a key to independently/voluntarily prove that they were the source of a given ring signature, which doesn't logically require culpability. Claimability is generally easy to achieve with some proof of knowledge technique.

The thoughtful reader probably won't have much trouble in imagining what it would mean to do the logical AND of 2 \(\Sigma\)-protocols.

"AND" here just means you need to prove to the Verifier that you know both secrets / both conditions are true. So this only requires that you can answer both challenges (second step) with correct responses. Using the standard notation, that means generating two transcripts:

\((R_1, e, s_1) \quad (R_2, e, s_2)\)

i.e. the same \(e\)-value is given to both protocol runs after receiving the initial commitments from each. Fiat-Shamir-ising this protocol will work the same as the usual logic; if considering a signature scheme, we'll be hashing something like \(H(m||R_1||R_2||P_1||P_2)\), if we include, as we have learnt to, key-prefixing.

As we already mentioned, the Borromean ring signature design uses this idea to compactify a set of ring signatures, since only one \(e\)-value is being published, rather than \(M\) for \(M\) ring signatures.

This much is not super-interesting; but we can tighten this up a bit and only use **one** commitment and response in a special case:

See one of the first posts on this blog for a description of this technique; here we're giving a slightly deeper look at the meaning.

If you are proving not only knowledge of a secret \(x\), but also that two curve points have the same discrete log \(x\) w.r.t. different bases \(G\) and \(J\) (whose relative discrete log must not be known; see earlier blog post etc.), you can condense the above AND by reusing the commitment and challenge for the two bases:

\(\mathbb{P} \rightarrow \mathbb{V}\): \(R_1= kG,R_2=kJ\)

\(\mathbb{V} \rightarrow \mathbb{P}\): \(e = H(m||R_1||R_2||P_1||P_2)\)

\(\mathbb{P} \rightarrow \mathbb{V}\): \(s\), (in secret: \(=k+ex\))

Now, if the prover acted honestly, his construction of \(s\) will correctly pass verification **twice**:

\(sG \stackrel{?}{=}R_1 +e P_1 \quad sJ \stackrel{?}{=} R_2 + eP_2\)

... and notice that it would be impossible to make that work for different \(x\)-values on the two bases \(G\) and \(J\) because you would need to find \(k_1, k_2 \in \mathbb{Z}_N, x_1, x_2 \in \mathbb{Z}_N\) such that, **without knowing \(e\) in advance, **\(s = k_1 + ex_1 =k_2 + ex_2\), which is clearly impossible.

Proof of soundness is easy to see using the standard rewinding technique (see e.g. previous blog post amongst many other places); after the two upfront commitments are fixed and the \(e\)-values are "forked", we will get two \(s\) values as usual and extract \(x\).

Shortly after the AOS paper, Liu, Wei and Wong published a paper outlining how the same basic idea could be extended to a slightly more complex context of requiring **linkability**, as earlier mentioned. It uses a combination of the above: DLEQ via AND of \(\Sigma\)-protocols, and OR of \(\Sigma\)-protocols for the ring signature hiding effect. Detailed algorithm with commentary follows.

We start with a keyset \(L = \{P_0 \ldots P_{n-1}\}\) chosen by the signer, whose index will be \(\pi\) (note the ambiguities about "what is the set of valid keys?" as was discussed under "Key Prefixing"). We then form a special new kind of curve point that we'll name from now on as the **key image** (for reasons that'll become clear):

\(I =x_{\pi} \mathbb{H}(L)\)

Here \(\mathbb{H}\) is a hash function whose output space is points on the curve, rather than scalar numbers. (*The mechanical operation for doing this is sometimes described as "coerce to point"; for example, take the 256 bit number output by SHA256 and interpret it as an \(x-\)coordinate on secp256k1, find the "next" valid point \(x,y\), incrementing \(x\) if necessary, or whatever; just has to be deterministic*). \(\mathbb{H}(L)\) is therefore going to play the same role as \(J\) in the previous section, and we assume intractability of relative discrete log due to the hashing.

The following steps are very similar "in spirit" to AOS; we still "extend the causality loop" (bastardising Poelstra's description) over the whole set of signatures instead of just one, but this time we also "lift" the loop onto a base of \(\mathbb{H}(L)\) and replicate the signatures there, too:

- Set \(k_{\pi} \stackrel{$}{\leftarrow} \mathbb{Z}_N\)
- Form the hash-challenge at the next index: \(e_{\pi+1} = H(m||L||k_{\pi}G||k_{\pi}\mathbb{H}(L)||I)\)
- Note to the above: \(k_{\pi}G\) was previously called \(R_{\pi}\) in AOS; we are trying to preserve here, the same notation where possible; and of course it's the \(R\) value, not the \(k\)-value that will be known/calculated by the verifier. The same applies to the "lifted" nonce-point which follows it in the concatenation. With respect to the key image, note that it
*will*be published and known to the verifier; but he won't know which index it corresponds to. - Pick \(s_{\pi+1} \stackrel{$}{\leftarrow} \mathbb{Z}_N\); then we do as in AOS, but duplicated; we set:
- \(R_{\pi+1} = s_{\pi+1}G - e_{\pi+1}P_{\pi+1}\) and \(R^{*}_{\pi+1} = s_{\pi+1}\mathbb{H}(L) - e_{\pi+1}I\)
- I realise the last line is pretty dense, so let's clarify: the first half is exactly as for AOS; calculate \(R\) given the random \(s\) and the just-calculated hash value \(e\). The
*second*half is**the same thing with the base point \(G\) replaced with \(\mathbb{H}(L)\), and the pubkey replaced with \(I\) at every index**. We used a shorthand \(R^{*}\) to mean \(k_{\pi+1}\mathbb{H}(L)\), because of course we don't actually*know*the value \(k_{\pi+1}\). - Calculate the next hash-challenge as \(e_{pi+2} = H(m||L||R_{\pi+1}||R^{*}_{\pi+1}||I)\)
- Etc...
- As with AOS, we can now forge all the remaining indices, wrapping around the loop, by repeating the above operation, generating a new random \(s\) at each step, until we get back to the signing index \(\pi\), when we must calculate \(s_{\pi}\) as: \(s_{\pi} = k_{\pi} + e_{\pi}x_{\pi}\).
- Signature is published as \(\sigma_{L}(m) = (s_0 \ldots s_{n-1}, e_0, I)\). (As before, if the keyset \(L\) is not specified in advance, it will have to be published for the verifier).

So what we're doing here is OR(DLEQ(0), DLEQ(1),.... DLEQ(n-1)). And as observed, each DLEQ is actually an AND: "AND(I know x for P, x for P is same as x for P2)". Hence this represents a clever combination of AND- and OR- of \(\Sigma\)-protocols.

*On a personal note, when I first saw something of this type (I think it was Cryptonote, see below), I found it quite bewildering, and I'm sure I'm not alone! But what partially saved me is having already studied PoDLE/DLEQ as well as AOS ring sigs, so I could intuit that something combining the two ideas was going on. I hope the previous painstaking introductions make it all a lot clearer!*

Note the key similarities and difference(s) in the published signature, to the AOS case: you still only need to publish one hash \(e_0\) since the others are determined by it, but you **must** publish also the key image \(I\); if another LWW LSAG is published using the same private key, it will perforce have the same key image, and be recognized as having come from the same key without revealing which key.

The protocol using the LSAG can thus reject a "double-sign", if desired.

Let's sanity check that we understand the verification algorithm, since it is slightly different than AOS:

Start with the given keyset \(L\), the message \(m\) and the signature \((s_0 \ldots s_{n-1}, e_0, I)\)

- Construct \(e_{1} = H(m||L||R_{0}||R^{*}_{0}||I)\) using \(R_0 = s_0G - e_0 P_0\) and \(R^{*}_{0} = s_0 \mathbb{H}(L) - e_0 I \)
- Repeat at each index using the new \(e_j\) until \(e_0\) is calculated at the last step and verify it matches: \(e_0 \stackrel{?}{=} H(m||L||R_{n-1}||R^{*}_{n-1}||I)\). Accept if so, reject if not.

(with the additional point mentioned: the protocol using the sig scheme may also reject this as valid if \(I\) has already been used; this additional protocol step is usually described as "LINK" in the literature).

Make sure you get the difference between this \(\mathbb{H}(L)\) and the previous \(J\) as per the general DLEQ. In the latter case we can (and should) choose an arbitrary globally-agreed NUMS point, for example hashing the standard curve base point \(G\) (with the "coerce-to-point" technique mentioned). In this case, we have chosen something that both signer and verifier agree on, as part of the **setting** of this particular run of the protocol - it's deterministically tied to the keyset \(L\). The key image\(I\) is analogous to \(P_2\) in my PoDLE blog post; it's the signer's "hidden", one-time key.

This changes in the next construction, Back 2015. But first, a few words on security.

The general approach to proving **unforgeability **of this ring signature is the same as that for the basic Schnorr signature as described in the previous blog post.

A wrapper around an attacker \(\mathbb{A}\) who we posit to have the ability to construct a forgery without knowing any private key \(x_i\), will, as before, have to guess which random oracle query corresponds to the forgery, and will want to provide two different "patched" answers to the RO query at that point. As before, there will be some reduced probability of success due to having to make this kind of guess, and so the reduction will be even less tight than before.

Also as before, in the EUF-CMA model, we must allow for an arbitrary number of signing oracle as well as RO queries, which complicates the statistical analysis considerably, but the basic principles remain the same. If at some point forgery is successfully achieved twice at the same index, we will have something like:

\(x_{\pi} = \frac{s^{*}_{\pi}-s_{\pi}}{e^{*}_{\pi}-e_{\pi}}\)

where the * superscripts indicate the second run, and the \(e\)-values being the patched RO responses.

And as usual, with appropriate statistical arguments, one can generate a reduction such that forgery ability with a certain probability \(p\) implies thus ability to solve ECDLP with a related probability \(p'\).

For proving **signer ambiguity **- for simplicity, we break this into two parts. If *all* of the private keys are known to the attacker (e.g. by subpoena), then this property completely fails. This is what we called **culpability**. It's easy to see why - we have the key image as part of the signature, and that is deterministically reproducible given the private key. If *none *of the private keys are known to the attacker, the problem is reduced to the **solution of the Decisional Diffie Hellman Problem**, which is considered computationally hard. The reduction is quite complicated, but as in a standard zero knowledgeness proof, the idea is that a Simulator can generate a transcript that's statistically indistinguishable from a genuine transcript.

For proving **linkability **- in the LWW paper an argument is made that this reduces to ECDLP in more or less the same was as for the unforgeability argument, using two pairs of transcripts for two different signatures which are posited to be based on the same private key but having different key images. Examination of the two pairs of transcripts allows one to deduce that the private key in the two cases are the same, else ECDLP is broken.

Notice that these security arguments are much more complicated than for the single Schnorr signature case and perhaps for two distinct reasons: one, because the ring signature is a more complex algebraic construction, with more degrees of freedom, but also, because we are asking for a significantly richer set of properties to hold. In particular notice that even for unforgeability, the EUF-CMA description is not good enough (we've already discussed this a bit); we need to consider what happens when creating multiple signatures on different keysets and how they overlap. Signer anonymity/ambiguity is especially difficult for LWW and its postdecessors (see below), because by design it has been weakened (culpability).

*This is a good place to note that the constructions starting with LWW are described in some detail in the useful document Zero-To-Monero*.

Adam Back posted in 2015 on bitcointalk about a potential space saving over the cryptonote ring signature, based on using AOS and tweaking it to include a key image.

As was noted above, it's a space saving of asymptotically about 50% to use a scheme like AOS that only requires publication of one hash challenge as opposed to one for each index (like the CDS for example).

He then followed up noting that a very similar algorithm had already been published, namely the LWW we've just described in the above, and moreover it was published three years before Fujisaki-Suzuki that was the basis of cryptonote (see below). So it was *somewhat* of an independent re-discovery, but there is a significant tweak. I'll outline the algorithm below; it'll look very similar to LWW LSAG, but there's a difference.

- Define key image \(I =x_{\pi}\mathbb{H}(P_{\pi})\);
- Set \(k_{\pi} \stackrel{$}{\leftarrow} \mathbb{Z}_N\)
- Form the hash-challenge at the next index: \(e_{\pi+1} = H(m||k_{\pi}G||k_{\pi}\mathbb{H}(P_{\pi}))\)
- Pick \(s_{\pi+1} \stackrel{$}{\leftarrow} \mathbb{Z}_N\); then:
- \(R_{\pi+1} = s_{\pi+1}G - e_{\pi+1}P_{\pi+1}\) and \(R^{*}_{\pi+1} = s_{\pi+1}\mathbb{H}(P_{\pi+1}) - e_{\pi+1}I\)
- Calculate the next hash-challenge as \(e_{pi+2} = H(m||R_{\pi+1}||R^{*}_{\pi+1})\)
- Etc...
- As with AOS and LWW, we can now forge all the remaining indices, wrapping around the loop, by repeating the above operation, generating a new random \(s\) at each step, until we get back to the signing index \(\pi\), when we must calculate \(s_{\pi}\) as: \(s_{\pi} = k_{\pi} + e_{\pi}x_{\pi}\).
- Signature is published as \(\sigma_{L}(m) = (s_0 \ldots s_{n-1}, e_0, I)\), as in LWW (\(L\) being the set of \(P\)s).

Verification for this is near-identical as for LWW, so is left as an exercise for the reader.

The tweak - which is very similar to Cryptonote (makes sense as it was an attempt to improve that) - is basically this: by making each of the signatures in the shifted base point version symmetrical (example: \(s_2 \mathbb{H}(P_2) = k_2 \mathbb{H}(P_2) + e_2 I\)), it means that a key image will be valid *independent of the set of public keys, \(L\).* This is crucial in a cryptocurrency application - we need the key image to be a unique double spend signifier across many different ring signatures with different keysets - the keys are ephemeral and change between transactions.

So it's a blend of the LWW LSAG, which has the advantage of space compaction for the same reason as AOS - only one hash must be published, the others can be deduced from the ring structure - with the F-S-2007/Cryptonote design, which fixes the key image to the key and not just the specific ring.

However I have to here leave open whether the security arguments of LWW carry across to this case. I note that the original description did* not *include the keyset in the hash challenge (notice absence of \(L\)); but see the note on MLSAG below.

Cryptonote was adapted from a paper of Fujisaki and Suzuki describing an alternate version of a linkable (here "traceable") ring signature, in 2007. We won't dwell on these constructions here (except inasmuch as we referred to them above), as they provide the same linkability function as the above LSAG, but are less compact. Instead, in the final section, I'll describe how Monero has applied LWW LSAG and the Back LSAG to their specific requirements.

For anyone paying close attention all the way through, there will be nothing surprising here!

For a cryptocurrency, we build transactions consisting of multiple inputs. Each input in Monero's case uses a ring signature, rather than a single signature, to authenticate the transfer, referring back to multiple pubkeys possessing coins as outputs of earlier transactions.

So here we need **one ring signature per input**. Moreover, per normal transaction logic, we obviously need *all* of those ring signatures to successfully verify. So this is another case for the "AND of \(\Sigma\)-protocols". We just run \(M\) cases of Back's LSAG and combine them with a single \(e\) hash challenge at each key index (so the hash challenge kind of "spans over the inputs"). Additionally, note that the hash challenge here is assumed to include the keyset with a generic \(L\) (limiting tiresome subscripting to a minimum...).

To sign \(M\) inputs each of which have \(n\) keys:

- For each input, define key image \(I_i =x_{i,\pi}\mathbb{H}(P_{i,\pi}) \ \forall i \in 0 \ldots M-1\);
- Set \(k_{i, \pi} \stackrel{$}{\leftarrow} \mathbb{Z}_N \ \forall i \in 0 \ldots M-1\)
- Form the hash-challenge at the next index: \(e_{\pi+1} = H(m||L||k_{0, \pi}G||k_{0, \pi}\mathbb{H}(P_{0,\pi})||k_{1, \pi}G||k_{1, \pi}\mathbb{H}(P_{1,\pi}) ...)\)
- Pick \(s_{i, \pi+1} \stackrel{$}{\leftarrow} \mathbb{Z}_N\ \forall i \in 0 \ldots M-1\); then:
- \(R_{i, \pi+1} = s_{i, \pi+1}G - e_{\pi+1}P_{i, \pi+1}\) and \(R^{*}_{i, \pi+1} = s_{i, \pi+1}\mathbb{H}(P_{i, \pi+1}) - e_{\pi+1}I_i \ \forall i \in 0 \ldots M-1\)
- Calculate the next hash-challenge as \(e_{\pi+2} = H(m||L||R_{0, \pi+1}||R^{*}_{0,\pi+1}||R_{1, \pi+1}||R^{*}_{2,\pi+1} ...)\)
- Etc...
- Logic as for AOS, LWW but duplicated at every input with single \(e\)-challenge, and at signing index for all inputs (\(\pi\)): \(s_{i, \pi} = k_{i, \pi} + e_{i, \pi}x_{i, \pi}\ \forall i \in 0 \ldots M-1\).
- Signature is published as \(\sigma_{L}(m) = (s_{0,0} \ldots s_{0,M-1}, \ldots, s_{n-1,0}, \ldots s_{n-1,M-1}, e_0, I_0 \ldots I_{M-1})\).

Note:

(1) This algorithm as described requires each input to have the genuine signer at the same key-index in the set of pubkeys for each input, which is a limitation.

(2) Monero has implemented Confidential Transactions, and this is folded in with the above into a new design which seems to have two variants "RingCTFull" and "RingCTSimple". This can be investigate further in the documents on RingCT as referenced in the previously mentioned ZeroToMonero.

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