Consider that I wish to prove knowledge of some RSA private key corresponding to a public key $(e,N)$. A naive interactive proof scheme would proceed as follows:

  • $V$ generates some random message $m$ and encrypts it, sending the encrypted data $c$ to $P$
  • $V$ then requests $m$ back from $P$. Assuming $P$ could have no knowledge of $m$ outside of $c$, then $P$ can prove knowledge of $d$ by responding with the correct $m$

This appears to be sound and complete on the assumption of hardness of the RSA problem. Clearly, this must be a private-coin system, since if random data was public then a dishonest $P$ would be able to deduce any $m$. I think we also need the assumption that this is an honest-verifier scheme on the basis that a decryption oracle would allow for a break of textbook RSA (I think). However, we could also abstract away the details of the asymmetric cryptosystem used, and assume it is some theoretical asymmetric cryptosystem which cannot be broken using a decryption oracle, and then I think we could drop this stipulation.

Intuitively, this scheme is not zero-knowledge, since we are giving away something somewhat significant (i.e. the decryption of an arbitrary $c$). However, we can easily construct a simulator for this scheme which will choose some random $m$, encrypt it, and then just send back the original $m$ - which implies that it is in fact zero-knowledge. Am I missing some nuance here with what the simulator can/cannot do?

Is this scheme zero-knowledge with RSA (why/why not)? If it isn't, would it be zero-knowledge with some idealised asymmetric scheme which is not vulnerable to any attacks based on decryption oracles (why/why not)?


1 Answer 1


I've done some reading since posting this question and I now have a better grasp on this, so I'll self-answer.

The confusion arises from the difference between zero-knowledge and honest-verifier zero knowledge. Consider the verifier $V^0$ which will generate a random $m$, encrypt it to generate a ciphertext $c$ and send it over the channel to $P$. $P$ (assumed to also be honest) will then respond with $m$. This interaction is trivially zero-knowledge. Intuitively, $V^0$ has learned nothing, since they already knew $m$ - formally we can build a simulator which creates a false transcript of the proof by generating a random $m$ and encrypting it. However, $V^0$ is the honest verifier - it is a verifier acting according to the protocol.

Consider the following malicious verifier $V^*$: it does not perform any encryptions and will simply always send a set ciphertext, $c^*$. Now something is learned - the decryption of $c^*$, which was not known before since $c^*$ is not the result of a known encryption, but the result of specific choice. The definition of zero-knowledge stipulates that for all verifiers there must exist a simulator. $V^*$ cannot be simulated without knowledge of the secret, since it is easy to check whether the decryption is correct by re-encrypting it, so the correct decryption of $c^*$ must be known by the simulator, in which case the simulator must be able to decrypt arbitrary ciphertexts, which is not possible without access to the secret; thus the scheme is in fact not zero-knowledge, regardless of the cryptosystem used (unless the cryptosystem is broken and therefore a simulator can decrypt arbitrary messages in polynomial time, in which case the proof is moot). Note here that it is essential $V^*$ sends a fixed ciphertext every time - if $c^*$ is picked at random for each interaction, the transcript can easily be falsified.

It's not a part of the original question, but I think it's important to bring up - this is not a valid proof-of-knowledge, which I intended it to be. In fact (denoting by $S_\mathcal{M}$ this scheme implemented with some cryptosystem $\mathcal{M}$), "$S_\mathcal{M}$ is a proof-of knowledge" $\implies$ "decryption oracle access to $\mathcal{M}$ allows private-key exfiltration". This is an immediate consequence of the lack of a commitment stage as would be used in a $\Sigma$-protocol. A proof-of-knowledge requires that we can write an extractor $E$ which can retrieve the secret from $P$ if allowed to rewind its state. However, $P$ as described here is stateless - regardless of which interaction or when in the interaction an input is sent, the output is the same - thus the existence of an efficient extractor for $P$ immediately implies that $\mathcal{M}$ is completely broken with decryption oracle access. This is merely an interactive proof that proves (with $\epsilon$, the margin for error, hazily defined) that $P$ can decrypt arbitrary messages - it does not satisfy the requirements to prove private key knowledge.


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