Abstract
The secure instantiation of the random oracle is one of the major open problems in modern cryptography. We investigate this problem using concepts and methods of algorithmic randomness. In modern cryptography, the random oracle model is widely used as an imaginary framework in which the security of a cryptographic scheme is discussed. In the random oracle model, the cryptographic hash function used in a cryptographic scheme is formulated as a random variable uniformly distributed over all possibility of the function, called the random oracle. The main result of this paper is to show that, for any secure signature scheme in the random oracle model, there exists a specific computable function which can instantiate the random oracle while keeping the security originally proved in the random oracle model. In modern cryptography the generic group model is used also for a similar purpose to the random oracle model. We show that the same results hold for the generic group model. In the process of proving the results, we introduce the notion of effective security, demonstrating the importance of this notion in modern cryptography.
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Notes
Normally, Martin-Löf random is defined with fixing the total recursive function \(f\colon \mathbb {N}^{+}\to \mathbb {Q}\cap (0,\infty )\) to the form f(n)=2 −n . However, the relaxation of the function f as in Definition 1 does not alter the class of Martin-Löf random infinite binary sequences.
Normally, a probabilistic (uniform) polynomial-time Turing machine is called a probabilistic polynomial-time adversary when it is used as an adversary against a cryptographic scheme.
In fact, \(\mathsf {ML\text {-}TEST}_{\mathrm {\Pi }}^{\mathsf {EUF\text {-}ACMA}}\) contains only Schnorr tests, where a Schnorr test is defined as a Martin-Löf test \(\mathcal {C}\subset \mathbb {N}^{+}\times \{0,1\}^{*}\) such that \(\mathcal {L}\left ({[{\mathcal {C}_{n}}]^{\prec }}\right )\) is computable uniformly in n. For the detail of Schnorr tests, see e.g. Section 3.5 of Nies [23].
Lemma 3 can be used to prove the non-existence of universal Schnorr test for the notion of Schnorr randomness for an infinite binary sequence. See Fact 3.5.9 of [23] for the detail.
See [15, Chapter 4] for the detail of collision resistant hash function.
As such S ′, the set T∪{s∈S∣t is a prefix of s} suffices, where T={t} if there is a prefix s∈S of t and T=∅ otherwise.
In order to prove Theorem 13, it is suffice to use the inequality 2n≥n d which holds for all n≥((d+1)/ ln2)d+1 and not for all n≥d 2 as in Lemma 4. The former follows immediately from the inequality e x≥x which holds for all \(x\in \mathbb {R}\). However, we prefer a more “insightful” polynomial lower bound n≥d 2 than the super-exponential lower bound n≥((d+1)/ ln2)d+1. See Section 11 for further remarks.
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This work was supported by JSPS KAKENHI Grant Number 23340020 and by the Ministry of Economy, Trade and Industry of Japan.
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A part of this work was presented at the Seventh International Conference on Computability, Complexity and Randomness (CCR 2012), July 2-6, 2012, Cambridge, Great Britain.
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Tadaki, K., Doi, N. Cryptography and Algorithmic Randomness. Theory Comput Syst 56, 544–580 (2015). https://doi.org/10.1007/s00224-014-9545-9
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DOI: https://doi.org/10.1007/s00224-014-9545-9