1 / 38

Computationally Sound Symbolic Protocol Analysis: Correspondence Theorems

18739A: Foundations of Security and Privacy. Computationally Sound Symbolic Protocol Analysis: Correspondence Theorems. Anupam Datta CMU Fall 2007-08. Protocol Analysis Techniques. Crypto Protocol Analysis. Formal Models. Dolev-Yao (perfect cryptography). Computational Models.

enya
Download Presentation

Computationally Sound Symbolic Protocol Analysis: Correspondence Theorems

An Image/Link below is provided (as is) to download presentation Download Policy: Content on the Website is provided to you AS IS for your information and personal use and may not be sold / licensed / shared on other websites without getting consent from its author. Content is provided to you AS IS for your information and personal use only. Download presentation by click this link. While downloading, if for some reason you are not able to download a presentation, the publisher may have deleted the file from their server. During download, if you can't get a presentation, the file might be deleted by the publisher.

E N D

Presentation Transcript

  1. 18739A: Foundations of Security and Privacy Computationally Sound Symbolic Protocol Analysis: Correspondence Theorems Anupam Datta CMU Fall 2007-08

  2. Protocol Analysis Techniques Crypto Protocol Analysis Formal Models Dolev-Yao (perfect cryptography) Computational Models Random oracle Probabilistic process calculi Probabilistic I/O automata • Correspondence theorems • Computational PCL … Model Checking Protocol Logics Process Calculi Inductive Proofs … Murphi, AVISPA BAN, PCL Applied -calculus Paulson, MSR

  3. Computational and Symbolic Analysis • Cryptography: computational constructions • Define actual algorithms operating on bit strings • For example, RSA is defined as a triple of algorithms • Key generation: public key is (n,e), private key is d where n=pq for some large primes p,q; ed=1 mod (p-1)(q-1) • Encryption of m: me mod n • Decryption of c: cd mod n • Formal model: abstract treatment of crypto • Instead of defining cryptographic algorithms, simply say that they satisfy a certain set of properties • Given ciphertext, obtain plaintext if have exactly the right key • Cannot learn anything from ciphertext without the right key

  4. Goal: Soundness of Formal Analysis • Prove that properties assumed in the formal model are true for some cryptographic constructions • Formal proofs must be sound in following sense: • For any attack in concrete (computational) model… • …there is matching attack in abstract (formal) model • …or else the concrete attack violates computational security of some cryptographic primitive • If we don’t find an attack in the formal model, then no computational attack exists • More precisely, probability that a computational attack exists is negligible

  5. Correspondence Theorems Abstract execution model Security properties Formal security Protocol specification language Soundness theorem Relation Concrete execution model Security properties Cryptographic security

  6. Representative Papers • Abadi, Rogaway. “Reconciling Two Views of Cryptography (The Computational Soundness of Formal Encryption)”. J. Cryptology, 2002. • Symmetric encryption only, passive adversary • Backes, Pfitzmann, Waidner. “Universally Composable Cryptographic Library”. CCS 2003. • Public key encryption, signatures; active adversary • Micciancio, Warinschi. “Soundness of Formal Encryption in the Presence of Active Adversaries”. TCC 2004. • Public-key encryption, active adversary Today

  7. Another Approach High-level proof principles PCL • Syntax (Properties) • Proof System (Proofs) Computational PCL • Syntax ±  • Proof System±  Soundness Theorem (Induction) Soundness Theorem (Reduction) [BPW, MW,…] Symbolic Model • PCL Semantics (Meaning of formulas) Cryptographic Model • PCL Semantics (Meaning of formulas) Unbounded # concurrent sessions Polynomial # concurrent sessions

  8. Representative Papers Next lecture • Datta, Derek, Mitchell, Shmatikov, Turuani. “Probabilistic Polynomial-time Semantics for a Protocol Security Logic”. ICALP 2005. • Public-key encryption, active adversary • A. Datta, A. Derek, J. C. Mitchell, B. Warinschi, Computationally Sound Compositional Logic for Key Exchange Protocols, in Proceedings of 19th IEEE Computer Security Foundations Workshop, pp. 321-334, July 2006. • Signatures, key exchange, secure sessions • A. Roy, A. Datta, A. Derek, J. C. Mitchell, Inductive Proofs of Computational Secrecy, in Proceedings of 12th European Symposium On Research In Computer Security , September 2007. • Symmetric and public key encryption; secrecy proof technique • A. Roy, A. Datta, J. C. Mitchell, Formal Proofs of Cryptographic Security of Diffie-Hellman based Protocols, in Proceedings of Symposium on Trustworthy Global Computing, November 2007. • Diffie-Hellman, symmetric and public key encryption, signatures.

  9. Outline • Define what it means for an encryption scheme to be secure against adaptive chosen-ciphertext attack in the multi-user setting • Define a formal language for describing protocols • Define concrete trace semantics for protocols • Actual execution traces of protocols obtained by instantiating nonces with bit strings, etc. • Traces include actions of the concrete adversary • Show that almost every action of concrete adversary maps to an action of abstract adversary • This will follow from security of encryption scheme

  10. Simple Protocol Language • Pairing and encryption only Term ::= Id | Identity Key | Public keys only Nonce | Pair | Ciphertext Pair ::= (Term, Term) Ciphertext ::= {Term}Key • Can write simple protocols with this syntax • Must describe valid computations of honest parties • For example, (B receives {X}pk(A), B sends {X}pk(B)) is not valid because B can’t decrypt {X}pk(A)

  11. Abstract vs. Concrete Execution Abstract execution Concrete execution (with RSA) A  {Na,A}pk(X) B  {Nb,NA}pk(A) A  {Nb}pk(X) A  bits(m1ex mod nx) B  bits(m2ea mod na) A  bits(m3ex mod nx) (ex, nx) is responder’s RSA public key; m1 is a number encoding a concatentation of a random bit string representing nonce Na and a fixed bit string representing A’s identity (ea, na) is A’s RSA public key; m2 is a number encoding a concatentation of a random bit string representing nonce Nb and the bit string extracted from the message received by B (ex, nx) is responder’s RSA public key; m3 is a number encoding the bit string extracted from the message received by A

  12. Abstract vs. Concrete Adversary Abstract adversary Concrete adversary Fix some set of corrupt users C. Let M be messages sent prior to some point in protocol execution. • M  know(C,M) • If t,t’  know(C,M), then (t,t’)  know(C,M) • If (t,t’)  know(C,M), then t,t’  know(C,M) • If t  know(C,M), then {t}k know(C,M) for any key k  Keys - Adversary can access any encryption oracle • If {t}ki know(C,M) and Ai C, then t  know(C,M) - Adversary can decrypt only messages encrypted with public keys of corrupt users Any polynomial-time algorithm (maybe probabilistic)

  13. Equivalence of Concrete & Abstract • Need to prove that concrete adversary cannot achieve more than the abstract adversary except with negligible probability • E.g., may guess secret key with negligible probability • Show that almost any concrete trace is an implementation of some valid abstract trace • Concrete traces represent everything that the concrete adversary can achieve • If almost any concrete trace can be achieved by the abstract adversary, it is sufficient to look only at the abstract adversary when doing analysis

  14. Concrete and Abstract Traces • Representation function R maps abstract symbols (nonces, keys, identities) to bit strings • Defines concrete “implementation” of abstract protocol • Concrete trace is an implementation of an abstract trace if exists a representation function R such that every message in concrete trace is a bit string instantiation of a message in abstract trace • Intuitively, concrete trace is an implementation if it is created by plugging bit strings in place of abstract symbols in the abstract trace • Denote implementation relation as 

  15. How Can This Fail? • Suppose encryption scheme is malleable • Can change encrypted message without decrypting • Plain old RSA is malleable! Abstract execution Concrete execution (with RSA) A  {Na,A}pk(B) B  ??? A  meb mod nb B  m2eb mod nb Adversary intercepts and squares; B receives m2eb instead of meb There is no abstract operation corresponding to the action of concrete adversary!

  16. Main Theorem • If the encryption scheme used in the protocol is IND-CCA secure, then for any concrete adversary Ac ProbRA,Ro(Af: traceset(Af, Of)  traceset(Ac(RA),Oc(RO))  1-() With overwhelming probability (i.e., dominated by any polynomial function of security parameter) Probability is taken over randomness of concrete adversary and encryption oracles (recall that concrete adversary may be probabilistic and public-key encryption must be probabilistic) Every concrete trace generated by concrete adversary interacting with concrete encryption and decryption oracles... (note that concrete adversary and concrete encryption oracles are randomized) There exists an abstract adversary such that… …some abstract trace generated by abstract adversary interacting with abstract encryption and decryption oracles is an implementation of…

  17. Proof Outline • For any (randomized) concrete adversary, construct the corresponding abstract adversary Af • Abstract “model” of concrete adversary’s behavior • Guarantees traceset(Af, Of)  traceset(Ac(RA),Oc(RO)) • Show that every action performed by constructed adversary is a valid action in the abstract model (with overwhelming probability) • Abstract adversary is only permitted to decrypt if he knows the right key (e.g., cannot gain partial info) • Prove: constructed adversary doesn’t do anything else

  18. Constructing Abstract Adversary • Fix coin tosses (i.e., randomness) of honest participants and concrete adversary • This uniquely determines keys and nonces of honest participants • Use symbolic constants to give them names in abstract model • Every bitstring which is not an honest participant’s key or nonce is considered the adversary’s nonce and given some symbolic name Must prove that every concrete action can be represented by a valid abstract action Concrete adversary Abstract adversary

  19. Reduction to IND-CCA • Prove that every time concrete adversary does something that’s not permitted in abstract model, he breaks IND-CCA security of encryption • This can only happen with negligible probability • Therefore, with overwhelming probability every concrete action implements some valid abstract action • We’ll prove this by constructing a simulator who runs the concrete adversary in a “box” and breaks IND-CCA exactly when the concrete adversary deviates from the abstract model

  20. CCA adversary interacts with encryption and decryption oracles according to rules of CCA game Overview of Reduction Encryption and decryption oracles (CCA game) Abstraction Af CCA adversary wins CCA game exactly when Af does something illegal in the abstract model CCA adversary CCA adversary runs concrete adversary as a subroutine To do this, CCA adversary must be able to simulate protocol to Ac (i.e., create an illusion for Ac that he is interacting with the actual protocol) Concrete adversary Ac

  21. What’s Illegal in Abstract Model? • The only way in which constructed abstract adversary Af may violate rules of abstract model is by sending a term with some honest nonce X that he could not have learned by abstract actions • X cannot be learned from messages sent by honest parties by simple decryption and unpairing rules • Case 1: If Af sends X in plaintext, this means that concrete adversary managed to open encryption • Recall that Af is abstraction of concrete adversary Ac • If concrete adversary extracts a bitstring (abstracted as nonce X) from under encryption, he wins the CCA game

  22. What if Encryption is Malleable? • Case 2: Af sends {t[X]}k, but neither t[X], nor {t[X]}k was previously sent by honest parties • Adversary managed to take some encryption containing X and convert it into another encryption containing X • This is known as malleability. For example, with RSA can convert an encryption of m into encryption of m2 • In this case, CCA adversary will win CCA game by using the concrete adversary’s ability to convert one encryption into another • But to do this, CCA adversary must be able to simulate protocol execution to the concrete adversary

  23. 2. Every time Ac asks for enck(s(X)), CCA adversary uses oracles to obtain enck(s(xb)) 3. When Ac outputs enck(t(xb)), CCA adversary submits it to decryption oracle 4. From t(xb) CCA adversary learns whether xo or x1 is inside; outputs bit b correctly s(xo), s(x1) enck(s(xb)) enck(s(xb)) enck(t(xb)) enck(t(xb)) t(xb) Ok, since the ciphertext was not previously created by encryption oracle Winning the CCA Game (Simplified) 1. CCA adversary guesses nonce X and picks two values xo and x1 CCA adversary b m0 m1 … Left-right selector LR(m0,m1,b) users Concrete adversary Ac

  24. Summary • If the encryption scheme used in the protocol is non-malleable under chosen-ciphertext attack, then the abstract model is sound • If an attack is not discovered in the abstract model, a concrete attack may exist only with negligible probability • See this paper for a similar result with a formal definition of the operational semantics of the simulator • A. Roy, A. Datta, A. Derek, J. C. Mitchell, Inductive Trace Properties for Computational Security, invited to appear in Journal of Computer Security.

  25. In more detail… • Abstract execution model and security properties • Concrete execution model and security properties • Relating abstract and concrete security properties

  26. I > R : {I,NI} (2) R > I : {R,NI,NR} (3) I > R : {NR} KR KI KR Semantics • A protocol defines two interactive programs, one for the Initiator and one for the Responder; • The program of a party A running the NSL protocol with intended responder B: - Choose random NA, encrypt (A,NA) and send the ciphertext to B; - Receive a message C; decrypt C with the secret key of A; parse the plaintext as (B, NA,NB); - If successful, send the encryption of NB under the public key of B.

  27. Semantics • The local state of a party: - program counter - an assignment from protocol variables to values, e.g. for the NSL protocol For the initiator A of a run between parties A, B the initial state: I=A, R=B, KI=KA,KR=KB,NI=NA, NR=? • The local state of the party changes by receiving/sending messages.

  28. A/B Party A new(A,B) A/B Party B M N Adversary can create new sessions of the protocol: new(A,B) Adversary can send messages to any party running the protocol and obtains the message the party would normally output Execution Model EnvironmentEnvmaintains the global state of all parties Adv

  29. I > R : {I,NI} (2) R > I : {R,NI,NR} (3) I > R : {NR} KR KI KR Abstract Execution Model EnvA maintains the global state of all parties; A/B Party A I=A ; R=B NI=NA ;NR=? A/B Party B I=A ; R=B NI=?;NR=NB new(A,B) 1 1 1 {A,NA} KB AdvA new(C,A) C/A Party C I=C ; R=A NI=NC ;NR=? C/A Party C I=C ; R=A NI=N;NR=NA 1 2 1 {C,NC} KA {C,N} KA Message are terms from a term algebra Adversary is a Dolev-Yao

  30. Abstract Execution Trace The interaction between each abstract adversary AdvAand the abstract environment EnvAdetermines an execution trace Denote by Tr(AdvA,EnvA)the trace F0,F1,F2,…,Fn, i.e. the sequence of the global state of the honest parties F0 F1 Fn …

  31. Abstract Security Properties An abstract security property is a predicate PropAon the set of all possible abstract execution traces A protocol satisfies abstractly the security property PropA if for all Dolev-Yao adversaries AdvATr(AdvA,EnvA) PropA

  32. I > R : {I,NI} (2) R > I : {R,NI,NR} (3) I > R : {NR} KR KI KR Concrete Execution Model EnvC maintains the global state of all parties a/b Party a I=a ; R=b NI=nA ;NR=? a/b Party b I=a ; R=b NI=?;NR=nB new(a,b) 1 1 {a,nA} 1 PKb AdvC new(c,a) c/a Party a I=c ; R=a NI=nC ;NR=? c/a Party c I=c ; R=a NI=n;NR=nA 1 2 1 {c,nc} PKa {c,n} PKa Messages are bit-strings Adversary is any probabilistic polynomial time Turing machine

  33. Concrete Execution Model Renv EnvC maintains the global state of all parties Radv a/b Party a I=a ; R=b NI=nA ;NR=? a/b Party b I=a ; R=b NI=?;NR=nB new(a,b) 1 1 {a,nA} 1 PKb AdvC new(c,a) c/a Party c I=c ; R=a NI=nC ;NR=? c/a Party c I=c ; R=a NI=n;NR=nA 1 2 1 {c,nc} PKa {c,n} PKa

  34. Concrete Execution Trace Fixing the randomness of the adversaries and of the parties, determines a unique execution trace Denote by Tr(AdvC(Radv),EnvC(Renv))the trace C0, C1, …, Cn determined by adversary AdvC using randomness Radvand environment EnvC using randomness Renv Renv Renv Renv Radv Radv Radv C0 C1 Cn …

  35. Concrete Security Properties A concrete security property is a predicate PropCon the set of all possible execution traces A protocol satisfies security property PropCif for any probabilistic polynomial-time adversaries AdvCTr(AdvC(Radv),EnvC(Renv)) PropCwith overwhelming probability over the random choices Radv and Renv

  36. Relating Abstract/Concrete Security Properties Fix PropA and PropC, an abstract and a concrete security property PropA < PropC if for all injective functions F, mapping formal values to appropriate bit-strings F(PropA)  PropC “Mutual authentication” can be naturally specified in both frameworks, and a relation as the one above holds

  37. Soundness Theorem THEOREM: Assume formal/computational security properties PropA<PropC and a protocol implemented with an IND-CCA secure encryption scheme. If the abstract execution of the protocol is secure (with respect to PropA), then the concrete execution is secure (with respect to PropC).

  38. Acknowledgements • Thanks to Vitaly Shmatikov and Bogdan Warinschi for providing a number of slides.

More Related