Abstract
Redactable signature schemes (\(\mathsf{RSS }\)) permit to remove parts from signed documents, while the signature remains valid. Some \(\mathsf{RSS }\)s for trees allow to redact non-leaves. Then, new edges have to be added to the tree to preserve it’s structure. This alters the position of the nodes’ children and may alter the semantic meaning encoded into the tree’s structure. We propose an extended security model, where the signer explicitly controls among which nodes new edges can be added. We present a provably secure construction based on accumulators with the enhanced notions of indistinguishability and strong one-wayness.
This is an extended and heavily revised version of [1]
The research leading to these results has received support from the European Union’s Seventh Framework Programme (FP7/2007–2013) under grant agreement no 609094.
Was supported by “Regionale Wettbewerbsfähigkeit und Beschäftigung”, Bayern, 2007–2013 (EFRE) as part of the SECBIT project (http://www.secbit.de) and the European Community’s Seventh Framework Programme through the EINS Network of Excellence under grant agreement no 288021, while at the University of Passau.
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Appendix
Appendix
1.1 Security Proofs of the Construction
We now show that our construction fulfills the given definitions. Namely, these are unforgeability, privacy, and transparency. We prove each property on its own.
Our Scheme is Unforgeable. If \(\mathcal {AH}\) is strongly one-way, while the signature scheme \(\varPi \) is unforgeable, our scheme is unforgeable.
Proof. Let \(\mathcal {A}\) be an algorithm winning the unforgeability game. We can then use \(\mathcal {A}\) in an algorithm \(\mathcal {B}\) to either to forge the underlying signature scheme \(\varPi \) or to break the strong one-wayness of \(\mathcal {AH}\). Given the game in Fig. 5 we can derive that a forgery must fall in at least one of the two following cases, for at least one node \(d\) in the tree:
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Type 1 Forgery: The value \(d\) protected by \(\sigma _s\) has never been signed by the signing oracle.
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Type 2 Forgery: The value \(d\) protected by \(\sigma _s\) has been signed, but \(T^* \notin \text {span}_\vdash (T, \sigma , \) \(\mathtt{ADM })\) for any tree \(T\) signed by the signing oracle.
Type 1 Forgery. In the first case, we can use the forgery generated by \(\mathcal {A}\) to create \(\mathcal {B}\) which forges a signature. We construct \(\mathcal {B}\) using \(\mathcal {A}\) as follows:
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1.
\(\mathcal {B}\) generates the key pair of \(\mathcal {AH}\), i.e., \(\mathtt{pk }\leftarrow \mathsf{KeyGen }(1^\lambda )\). It passes \(\mathtt{pk }\) to \(\mathcal {A}\). This is also true for \(\mathtt{pk }_S\) of the signature scheme to forge.
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2.
All queries to the signing oracle from \(\mathcal {A}\) are genuinely answered with one exception: instead of signing digests itself, \(\mathcal {B}\) asks it own signing oracle to generate the signature. Afterward, \(\mathcal {B}\) returns the signature generated to \(\mathcal {A}\).
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3.
Eventually, \(\mathcal {A}\) outputs a pair \((T^*, \sigma ^*)\). \(\mathcal {B}\) looks for the message/signature pair \((m^*,\) \(\sigma _s^*)\) inside the transcript not queried to its own signing oracle, i.e., the accumulator value with the signature \(\sigma _s^*\) of the root of \((T^*, \sigma ^*)\). Hence, there exists a value not signed by \(\mathcal {B}\)’s signing oracle. This pair is then returned as \(\mathcal {B}\)’s own forgery attempt.
As every tree/signature pair was accepted as valid, but not signed by the signing oracle, \(\mathcal {B}\) breaks the unforgeability of the signature algorithm. Here, we have a tight reduction for the first case.
Type 2 Forgery. In the case of a type 2 forgery, we can use \(\mathcal {A}\) to construct \(\mathcal {B}\), which breaks the strong one-wayness of the underlying accumulator. We construct \(\mathcal {B}\) using \(\mathcal {A}\) as follows:
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1.
\(\mathcal {B}\) generates a key pair of a signature scheme \(\varPi \).
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2.
It receives \(\mathtt{pk }\) of \(\mathcal {AH}\). Both public keys are forwarded to \(\mathcal {A}\).
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3.
For every request to the signing oracle, \(\mathcal {B}\) uses its hashing oracle to generate the witnesses and the accumulators. All other steps are genuinely performed. The signature is returned to \(\mathcal {A}\).
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4.
Eventually, \(\mathcal {A}\) outputs \((T^*, \sigma ^*)\). Given the transcript of the simulation, \(\mathcal {A}\) searches for a pair \((w^*,y^*)\) matching an accumulator \(a\), while \(y^*\) has not been queried to hashing oracle under \(a\). Note, the root accumulator has been returned: otherwise, we have a type 1 forgery. \(\mathcal {B}\) outputs \((a, w^*,y^*)\).
As every new element accepted as being part of the accumulator, while not been hashed by the hashing oracle, breaks the strong one-wayness of the accumulator, we have a tight reduction again.
Our Scheme is Private. If \(\mathcal {AH}\) is indistinguishable our scheme is private. Note: the random numbers do not leak any information, as they are distributed uniformly and are not ordered. Hence, we do not need to take them into account.
Proof. Let \(\mathcal {A}\) be an algorithm winning the privacy game. We can then use \(\mathcal {A}\) in an algorithm \(\mathcal {B}\) to break the indistinguishability of the accumulator \(\mathcal {AH}\). We construct \(\mathcal {B}\) using \(\mathcal {A}\) as follows:
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1.
\(\mathcal {B}\) generates a key pair of a signature scheme \(\varPi \).
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2.
It receives \(\mathtt{pk }\) of \(\mathcal {AH}\). Both public keys are forwarded to \(\mathcal {A}\).
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3.
For every request to the signing oracle, \(\mathcal {B}\) produces the expanded trees given \(\mathtt{ADM }\). Then, it uses its hashing-oracle to generate the accumulators, and then proceeds honestly as the original algorithm would do. Finally, it returns the generated signature \(\sigma \) to \(\mathcal {A}\).
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4.
For queries to the Left-or-Right oracle, \(\mathcal {B}\) extracts the common elements to be accumulated for both trees — this set is denoted \(\mathcal {S}\). Note, \(\mathcal {S}\) may be empty. The additional elements for the first hash are denoted \(\mathcal {R}_0\), and \(\mathcal {R}_1\) for the second one. \(\mathcal {B}\) now queries its own Left-or-Right oracle with \((\mathcal {S},\mathcal {R}_0,\mathcal {R}_1)\) for each hash. The result is used as the accumulator and the witnesses required: \(\mathcal {B}\) genuinely performs the rest of the signing algorithm and hands over the result to \(\mathcal {A}\).
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5.
Eventually, \(\mathcal {A}\) outputs its own guess \(d\).
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6.
\(\mathcal {B}\) outputs \(d\) as its own guess.
As we only pass queries, \(\mathcal {B}\) succeeds, whenever \(\mathcal {A}\) succeeds.
Our Construction is Transparent. We already know that our scheme is private. As neither the underlying signature, nor the witness’ values, nor the accumulator itself change during a redaction, no building block leaks additional information. Transparency follows.
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de Meer, H., Pöhls, H.C., Posegga, J., Samelin, K. (2014). Redactable Signature Schemes for Trees with Signer-Controlled Non-Leaf-Redactions. In: Obaidat, M., Filipe, J. (eds) E-Business and Telecommunications. ICETE 2012. Communications in Computer and Information Science, vol 455. Springer, Berlin, Heidelberg. https://doi.org/10.1007/978-3-662-44791-8_10
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DOI: https://doi.org/10.1007/978-3-662-44791-8_10
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