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Fully Homomorphic Encryption without Squashing Using Depth-3 Arithmetic Circuits Craig Gentry and Shai Halevi IBM T.J. Watson Research Center September 14, 2011 Abstract We describe a new approach for constructing fully homomorphic encryption (FHE) schemes. Previous FHE schemes all use the same blueprint from [Gentry 2009]: First construct a somewhat homomorphic encryption (SWHE) scheme, next “squash” the decryption circuit until it is simple enough to be handled within the homomorphic capacity of the SWHE scheme, and finally “bootstrap” to get a FHE scheme. In all existing schemes, the squashing technique induces an additional assumption: that the sparse subset sum problem (SSSP) is hard. Our new approach constructs FHE as a hybrid of a SWHE and a multiplicatively homomorphic encryption (MHE) scheme, such as Elgamal. Our construction eliminates the need for the squashing step, and thereby also removes the need to assume the SSSP is hard. We describe a few concrete instantiations of the new method, including a “simple” FHE scheme where we replace SSSP with Decision Diffie-Hellman, an optimization of the simple scheme that let us “compress” the FHE ciphertext into a single Elgamal ciphertext(!), and a scheme whose security can be (quantumly) reduced to the approximate ideal-SIVP. We stress that the new approach still relies on bootstrapping, but it shows how to bootstrap without having to “squash” the decryption circuit. The main technique is to express the decryption function of SWHE schemes as a depth-3 ( ∑ ∏ ∑ ) arithmetic circuit of a particular form. When evaluating this circuit homomorphically (as needed for bootstrapping), we temporarily switch to a MHE scheme, such as Elgamal, to handle the ∏ part. Due to the special form of the circuit, the switch to the MHE scheme can be done without having to evaluate anything homomorphically. We then translate the result back to the SWHE scheme by homomorphically evaluating the decryption function of the MHE scheme. Using our method, the SWHE scheme only needs to be capable of evaluating the MHE scheme’s decryption function, not its own decryption function. We thereby avoid the circularity that necessitated squashing in the original blueprint. Key words. Arithmetic Circuits, Depth-3 Circuits, Homomorphic Encryption, Symmetric Polynomials 1. Fully Homomorphic Encryption without Squashing
Contents 1 Introduction 1 1.1 Our Main Technical Innovation . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 1 1.2 An Illustration of an Elgamal-Based Instantiation . . . . . . . . . . . . . . . . . . . . 2 1.3 Leveled FHE Based on Worst-Case Hardness . . . . . . . . . . . . . . . . . . . . . . 3 2 Decryption as a Depth-3 Arithmetic Circuit 4 2.1 Restricted Depth-3 Arithmetic Circuits . . . . . . . . . . . . . . . . . . . . . . . . . 4 2.2 Lattice-Based Somewhat-Homomorphic Cryptosystems . . . . . . . . . . . . . . . . . 6 2.3 Decryption Using a Restricted Depth-3 Circuit . . . . . . . . . . . . . . . . . . . . . 6 3 Leveled FHE from SWHE and MHE 7 3.1 Notations . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 7 3.2 Compatible SWHE and MHE Schemes . . . . . . . . . . . . . . . . . . . . . . . . . . 8 3.3 Chimeric Leveled FHE: The Construction . . . . . . . . . . . . . . . . . . . . . . . . 9 4 Optimizations 10 4.1 Computing Only One Product . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 10 4.2 Short FHE Ciphertexts: Decryption as a Pure Symmetric Polynomial . . . . . . . . 11 References 12 A Instantiations of Chimeric FHE 14 A.1 The Homomorphic Capacity of SWHE Schemes . . . . . . . . . . . . . . . . . . . . . 14 A.2 Elgamal-based Instantiation . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 14 A.3 Leveled FHE Based on Worst-Case Hardness . . . . . . . . . . . . . . . . . . . . . . 16 A.3.1 Decryption under Sml . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 17 A.3.2 The SWHE scheme Lrg. . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 18 A.3.3 Setting the parameters. . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 18 B Proof of Lemma 1 19 1. Fully Homomorphic Encryption without Squashing
Page 1 and 2: AFRL-RI-RS-TR-2013-191 ADVANCED HOM
Page 3 and 4: REPORT DOCUMENTATION PAGE Form Appr
Page 5 and 6: 12. An Equational Approach to Secur
Page 7 and 8: 1.0 Executive Summary The ultimate
Page 9 and 10: The PROCEED program efforts are bro
Page 11 and 12: Towards the end of the project we i
Page 13 and 14: feature of the framework is that it
Page 15 and 16: 4.3 Publications Figure 1: A block
Page 17 and 18: encrypted data. We introduce a prec
Page 19 and 20: present an attack showing that a pa
Page 21 and 22: alternatives: mis (based on mutual
Page 23 and 24: 5.0 Conclusions and Recommendations
Page 25 and 26: [14] Bogdanov, Andrej, and Lee, Chi
Page 27 and 28: 8.0 List of Symbols, Abbreviations,
Page 29: Public Affairs Approvals for papers
Page 33 and 34: t j = P (a j ), and finally interpo
Page 35 and 36: apparent problem by replacing the M
Page 37 and 38: As noted above, there is a multilin
Page 39 and 40: f(x 1 , . . . , x n ) ∈ F and eve
Page 41 and 42: Translation. Pick up from the publi
Page 43 and 44: Lemma 4. Let prime p = ω(N 2 ). Th
Page 45 and 46: [vDGHV10] Marten van Dijk, Craig Ge
Page 47 and 48: have u 2 − v 2 = 1 → (u − v)(
Page 49 and 50: inary), and we want to compute g c
Page 51 and 52: Fully Homomorphic Encryption withou
Page 53 and 54: scheme have bit length Ω(D) to all
Page 55 and 56: fast. Thus, the actual magnitude of
Page 57 and 58: Definition 3 (B-bounded distributio
Page 59 and 60: achieve 2 λ security against known
Page 61 and 62: Proof. 〈c 2 , s 2 〉 = BitDecomp
Page 63 and 64: • FHE.Add(pk, c 1 , c 2 ): Takes
Page 65 and 66: The computation needed to compute t
Page 67 and 68: If a and b are large enough, B domi
Page 69 and 70: 5.1.2 How Batching Works Deploying
Page 71 and 72: We have the following lemma. Lemma
Page 73 and 74: Unpack(c, {τ σi→1 (s 1 )→s 2
Page 75 and 76: Since ‖e q1 ‖ < r q1 ,in, e q1
Page 77 and 78: [22] Damien Stehlé and Ron Steinfe
Page 79 and 80: 1 Introduction Fully homomorphic en
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encryptions of the same message usi
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and Tromer [CT10] in a completely d
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is negligible in k, where Expt CPA
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2.3 Non-Interactive Simulation-Soun
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is defined as (state 1 , m 1 , . .
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scheme has almost-all-keys perfect
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The algorithm S 1 . The algorithm S
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The distribution D 6 . This distrib
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We resolve this issue using the app
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the number of common reference stri
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• Homomorphic evaluation: The hom
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The number of repeated homomorphic
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[Gro04] [Gro10] J. Groth. Rerandomi
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Trapdoors for Lattices: Simpler, Ti
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mentioned, two central objects are
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Dimension m Quality s Length β ≈
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defined by G (or the semi-random ma
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1.4 Other Related Work Concrete par
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Cryptographic problems. For β > 0,
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andom over G m s , for uniformly ra
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3 Search to Decision Reduction Here
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• Preimage sampling for f G (x) =
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Note that for x ∈ {0, . . . , q
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The offline ‘bucketing’ approac
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G is primitive, the tag H in the ab
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conditional distribution of R given
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and parallelism of Algorithm 3 come
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is a coset of the lattice Λ = L( [
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scheme is su-scma-secure if Adv su-
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a discrete Gaussian with parameter
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• G ∈ Z n×nk q is a gadget mat
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must return this e (but possibly di
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[Gen09b] C. Gentry. Fully homomorph
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[vGHV10] M. van Dijk, C. Gentry, S.
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Contents 1 Introduction 1 2 Backgro
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1 Introduction In his breakthrough
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Then in Section 3 we describe our
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In some more detail, the BGV key sw
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itself has low norm. In BGV [5] thi
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3.4 Randomized Multiplication by Co
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One subtle implementation detail to
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ciphertexts. Producing the encrypte
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[24] Benny Pinkas, Thomas Schneider
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A.4 Sampling From A q At various po
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To maintain the noise estimate, the
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variance σ 2 φ(m), so 2d 2 (ζ)ɛ
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We stress that the only place where
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C.1.1 LWE with Sparse Key The analy
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The Smallest Modulus. After evaluat
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down by a factor of p t . Let’s r
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to apply a map κ i we express it a
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1 Introduction Fully homomorphic en
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4. No restrictions on the secret ke
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We point out that an alternative so
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Recall that GapSVP γ is the (promi
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Proof. By definition ⟨c, (1, s)
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• SI-HE.Enc pk (m): Identical to
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Proof of Theorem 4.2. Consider the
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We can now plug in what we know abo
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In particular (recall that E < ⌊q
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[Reg03] Oded Regev. New lattice bas
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1 Introduction An encryption scheme
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need to have errors. Since the expe
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Again we allow some “slackness”
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Learning Linear Functions. the clas
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Let us first outline the proof of T
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P ′ . 4 Namely H n ′ :n = P ′
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In conclusion, we get a sub-scheme
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[BV11b] [Gen09] [Gen10] [GH11] [GHS
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Quantum-Secure Message Authenticati
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the success probability away from 1
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Functions will be denoted by capito
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Quantum chosen message queries. In
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Since the w are the possible result
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number of distinct component values
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4.1 A Tight Upper Bound Theorem 4.1
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The algorithm is similar to the alg
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As we have already seen, for expone
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where S i (m) = (R i (m), PRF(k, (R
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To prove this theorem, we treat Y a
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Now we use this attack to obtain an
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Using this lemma we show that (3q +
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[KK11] Daniel M. Kane and Samuel A.
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1 Introduction Cloud storage system
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subset of the server’s storage, t
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security does not suffice. The main
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Static Data. PoR and PDP schemes fo
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For any efficient adversarial serve
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Overview of Construction. Let (Enc,
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means that one of the audit protoco
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Now we claim that an execution of E
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having higher capacity. The tables
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OWrite(i 1 , . . . , i t ; v 1 , .
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j-th level table again. With this o
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and distribute reprints for Governm
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A Simple Dynamic PoR with Square-Ro
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from Section 6 that is NRPH secure.
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The LWE problem was introduced in t
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1.2 Techniques and Comparison to Re
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(large) uniform errors as in [10],
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Proof. Let A be an algorithm runnin
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that 1 + ɛ = (1 + ɛ 1 )(1 + ɛ 2
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Because we are concerned with small
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For each i, since gcd(z i , q) = 1
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Proof. Let ω n = ω( √ log n) sa
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Theorem 3.8. Let m, n be integers,
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σ · O( √ m), except with probab
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k = n/2, and it suffices to set the
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[23] C. Peikert and B. Waters. Loss
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1 Introduction Consider the followi
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In the online phase clients individ
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the clients jointly run a secure co
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to compute F from scratch. The work
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I. Pre-processing phase. In this ph
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Online Phase: 1. Each client D i ge
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2n ciphertexts in order to be sure
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[KMR11] [KMR12] [Lip12] [LTV12] Sen
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For any set of adversarial parties
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C.5 Secure Computation We make use
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Experiment H 4 . This experiment is
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of A only in the case when A guesse
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leading to concrete implementations
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y modeling each block by an interac
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z i P 2 Q 1 y i P 1 s 1 r 1 x i w i
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2 Distributed Systems, Composition,
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[G | H](y) = (w, x): z = y w = y x[
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infinite chains, such as (M ∗ ;
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2.3 Multi-party computation, securi
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BCast(x) = (y 1 , . . . , y n ): y
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Sim(y A , s A ) = (y ′ A , r A):
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WCast(x ′ , w) = (y 1 ′ , . . .
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s ′ 0 s ′ A r ′ A y ′ H s
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p r ′ A Net’ s ′ r ′ H r A
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Notice that we have already underta
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simpler protocol description. For t
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[29] Andrew Chi-Chih Yao. Protocols
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2 Mihir Bellare, Stefano Tessaro, a
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10 Mihir Bellare, Stefano Tessaro,
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Multi-Instance Security and its App
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Multi-instance Security and Passwor
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a hash-of-hash construction: H 2 (M
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guisher queries (our lower bounds r
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in so doing they accidentally intro
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of D is defined as [ [ ] AdvH[P],R,
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main POW H[P],n,l1 : P 1 P 2 X 2
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now,thinkforconvenienceofw = l).The
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aroundO(q 2 )queries.Ideally, wewou
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HMAC l (K,Y 0) Y 0 F K Y 0 ′ G K
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This material is based on research
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24. Ari Juels and John G. Brainard.
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Contents 1 The BGV Homomorphic Encr
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‖a‖ canon 2 . The basic operati
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EncryptedArray/EncrytedArrayMod2r R
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g i ’s and h i ’s in one list,
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2.6 IndexSet and IndexMap: Sets and
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turn, are sometimes partitioned fur
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DoubleCRT& operator-=(const ZZ &num
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homomorphic automorphism operation
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For reasons that will be discussed
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where D i is the size of the i’th
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When key-switching a ciphertext ⃗
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constant (i.e. |poly(τ m )| 2 ), o
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ool isCorrect() const; and double l
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For the noise estimate in the new c
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The first time that ImportSecKey is
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EncryptedArray(const FHEcontext& co
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The MUX operation is implemented by
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5.1 Homomorphic Operations over GF
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EncryptedArrayMod2r ea(context); lo
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H/2m each and 0 with probability 1
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So now consider the inner sum in (7
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