C\s WSf¼Ê Kå ó6'"ü5c "üf> ] ±ESIGN ¼Ê². } Äntê.0 }z ep.0 1 ¹ã T R;<8 JXQ;< 8eph6 &u z2ÿå.056 y Cå +Y ¹ãC:G ÎP. ESIGN ¼Ê âå *Ð ûzôpö 4^ } ~ 2!


 Georgiana Terry
 1 years ago
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1 1 B=>? ESIGN 5789:46 C\s WSf¼Ê Kå ó6'"ü5c "üf> ] ±ESIGN ¼Ê²!mÖ} Ö¾ bï.0 z mš } Äntê.0 }z ep.0 1 ¹ã T R;<8 JXQ;< 8eph6 &u z2ÿå.056 y Cå +Y ¹ãC:G ÎP ÍY 5Õ Ì.0Ÿa ESIGN ¼Ê âå *Ð ûzôpö 4^ } ~ 2!mÖÊI z ƒ y $!mö Ö¾.0 z ˆ mš 2 ; < 2.1 Ž Š Ÿ WSf¼Ê 2 'E!mÖ ¾ å ö! Ÿ / 2¾ y {!mö ã Feige, Fiat, Shmair Q ŠŸ 2Ÿ,Â å y [8] $ ˆ ã Q Ÿ 2.1 ¼Ê5c (A; B) A: ÊIfª B: $Êf }z ª'Ú hã G ª A ó6' x )F' w z Ú ª x $Êf éw «ƒƒ ª' C: G k WNQ «(A; B) $ Ñ# Ds!m y z{ 1. (A(w);B)(x) ª ëåp > 1 0 (k) B in 2. $ { ö!f Adv ªu3 z ö!f Adv $Êf (A(w); Adv)(x) wúc5õeù ªˆ ðªêif B z ¼ÊZfQAd (Adv; B)(x) eù ƒ ª (Adv; B)(x) }z B in p G[ z > (k) Ì p ª G, Adv, B VÕ ã 2.2 ESIGN Ž Š ESIGN¼Ê { P9EKdÆH WS ŠWSf¼Ê Ïý ` ãn q ƒ ãn ÊI ã œbi y 1
2 Ÿ 2.2 ESIGNÆHæ,åk~2Çö! z u3å q/2 y ES IGN ¼Ê Ì ã 'E ¾ å ö! Ÿ /2¾ z{'e!mö y $Ì ãn ESIGNÆH!mÖ Ñ# ` ƒ í Ÿ 2.3 ESIGN ¼Ê e Í1]/ã 2.4 i ã a Ù z z{ñ# bgl]=b?d_pd }z ¾ å ö! Ÿ /2¾ z{'e!mö y $ ESIGN¼Ê!mÖ 1 ESIGNÆH!mÖ z mš 2.3 ESIGN ª Š!mÖ z ESIGNÆH RSA/ã ]Á y eí1]/ã (AERP) bgl]=b?d_pd  2 'E!m æ,åk~2çö! z u3å q/2 y ƒ ÊI (!mö z ` i Ÿ{È Ë mš ƒ!mö z ESIGN ÆH v {å!mö ÊI zš5c RSA Ž ÆH5c v PSS FDHRSA, } x*j Ž ÆH5c v ECSchnorr ( Š a  1:!mÖ ( 5c!mÖ ØÕxå bgl]õ 2 'E /ã /ã ESIGN!mÖÊI AERP Ó bgl] PSS or FDHRSA!mÖÊI RSA Ó bgl] ECSchnorr!mÖÊI x*oxzõ Ó bgl] 2.4 ESIGN ªŠ ƒ i ESIGN ÆH nxå!mö.0 a } nxå ÊI ù Œ Šx2 [9] C\s 2ð ç0 } ƒ x2 C\s ESIGNÆH TSH ESIGN É z ƒ ß' 2
3 Ÿ 2.4 G ESIGN 'Ú e Í1]L} (AERP) pk := fn; eg Gen(1 k ) y R f0; 1g k01 Q Š 0jjy =[x e mod n] k { x 2 (Z=nZ)npZ & L} y { påwúc_8dbcg] Adv z Ÿ m ãõ c, j1 Ü k z Pr[Adv(k; n; e; y)! x] < 1=k c q e Í1]L}» z z{ ƒƒ 0jjy =[x e mod n] k y p G Adv pì g z e Í1]L}» z z{/ã e Í1]/ã î Ÿ 2.5 ESIGN ÆH e Í1]/ãÙ z z{ñ# bgl]=b?d_ Pd }z æ,åk~2çö! z u3å q/2 y ± 1: e Í1]/ã z ºb CÆHÕ "üzc\o9y y C ESIGN ÆHÕ 15 ½l À Š [14, 13] ˆ $U ƒ C ESIGN ÆHÕ Cå!mÖ (5õÖ y e Í1]/ã z R %ù Š Õ `Õ e 4 $Ì Ïý z ì3 Nô ö! À& } [5, 6, 10, 18] ä#f n)õ14$7 Nô ö!6gz Oo z ± 2: n = p 2 q n)õ14 0»Ö z n = p 2 q n)õ14 n = pq ŸÛ~ {~ z 1~Œ z n = p 2 q ¹. Šz ~ 8dBcG ]% z [15, 16, 17, 1] ~ ƒ 8dBcG] z Ÿn)Õ14 x*j6 8dBcG] y (5 n = pq n = p 2 q z }z Ÿ2 ût n)õ148dbcg]õy z6 y ƒ 8dBcG] eù_ n C:G &u n)õ C:G &u z $Ì ì_è n = p 2 q C:G n = pq C:G!mÖ Ÿ Ÿ ø 3 < 3.1 ÆÑÅ µ Óœ Œ Ò ÈÏ ½: Hd[ F š : VerilogXL ³ DesignCompiler 3
4 ÎÀÑ½Ò : M 25.6KG(µ2NAND JhY) ³^_c (13312bit) [bgl]fen?³í Y Õ³1Y] ³ [^_c (13312bit)] ž Œ:?fN? 30MHz +Ytê (D\`e Dag rã) ÊI ÆH 9.8 ms $Ê 2.5 ms Š 'C:G 1152 bit 3.2 ÀÇÄ µ Óœ Œ ŒÈÍÂÄÉÑÊ: CPU: Pentium with MMX 266MHz OS: Trubo Linux version 4.0 ÔØÙ: íñ (gcc version ) w âøõ+yb:ybc gnu mp (gmp version 3.0.1) ZS ËÌÎÒ (»ÑÅ ): ËÌÎÒ (ÐÑº¹Îµ): ÊI ÆH Kbytes $Ê Kbytes ž : ÊI ÆH 512 Kbytes $Ê 476 Kbytes ÊI ÆH 18.0 ms $Ê 2.44 ms 4
5 Š 'C:G 1152 bit ÃÑÁ¼ ¾: n C:G 1152 bits e Ü 1024 hlen glen 32â ó6'x7:dc:g )F'X7:dC:G ÆH2X7:dC:G 160 bits 160 bits 128 bits 329 bytes 206 bytes 290 bytes Ú «: AgV:d_ 2æ. gcc O3 SzŠ C.0 (call6) ç0 ŠŸ eù Š"3 y mod n mod p 2 +Y }z ÞþÔ ÎPãn WS ût +Y ƒ ku ÿ5.0 ƒ ût.56 Sz z z ŠŒ M` g ù C.0 ût ep2¾ y 4 ÛAÝƒ Þ Ü ESIGN¼Ê ¹á ˆ ûôp y IfÝdÊI¼Ê5c FiatShamir ¼Ê ÖòÒ¼Ê5c x* Schnorr ¼Ê Ÿôptz ` ESIGN ôpö v {å 5c FiatShamir ¼Ê x* Schnorr ¼Ê ( ˆ 5c z 1152 WNQÕ ÎPÍY+T 5Õ.0 ƒ Œ 'Ú $Ê +T Änu &hÿœš ƒ ( }z 9 à U:Sc6 ÞþÔÎPãn,Âl ž ør.0 ù ŒŠ Vb^ K ESIGN FiatShamir ¼Ê5c 6 n 1152 WNQ x* Schnorr 5c %Õ 160 WNQ /ã Š Š, ESIGN e =2 5 FiatShamir ò Ò5Õ!m Õ 30 5 /ã Š } $  #M(1152) 1152 bits 6  ÎPÍY 5Õ 'E 8Änu #M(1152) Y ŠÜ a z 5
6  2: Änu ( 5c ÊIÄn $ÊÄn #M(1152) #M(1152) òò5õ (5) ESIGN x* Schnorr FiatShamir ²ßˆà [1] Adleman, L.M. and McCurley, K.S.: Open Problems in Number Theoretic Complexity,II (open problems: C7, O7a and O7b), Proc. of ANTSI, LNCS 877, SpringerVerlag, pp (1995). [2] Bellare, M. and Rogaway, P.: Random Oracles are Practical: A Paradigm for Designing Ecient Protocols, Proc. of the First ACM Conference on Computer and Communications Security, pp.62{73 (1993). [3] Bellare, M. and Rogaway, P. : Optimal Asymmetric Encryption, Proc. of Eurocrypt'94, LNCS 950, SpringerVerlag pp (1995). [4] Bellare, M. and Rogaway, P.: The Exact Security of Digital Signatures { How to Sign with RSA and Rabin, Proc. of Eurocrypt'96, LNCS 1070, SpringerVerlag, pp (1996). [5] Brickell, E. and DeLaurentis, J.: An Attack on a Signature Scheme Proposed by Okamoto and Shiraishi, Proc. of Crypto'85, LNCS 218, SpringerVerlag, pp (1986). [6] Brickell, E. and Odlyzko: Cryptanalysis: A Survey of Recent Results, Chap.10, Contemporary Cryptology, Simmons (Ed.), IEEE Press, pp.501{540 (1991). [7] Canetti, R., Goldreich, O. and Halevi, S.: The Random Oracle Methodology, Revisited, Proc. of STOC, ACM Press, pp.209{218 (1998). 6
7 [8] Feige, U., Fiat, A. and Shamir, A.: ZeroKnowledge Proofs of Identity, Journal of Cryptology, 1, 2, pp (1988) [9] Fujisaki, E. and Okamoto, T.: Security of Ecient Digital Signature Scheme TSHESIGN, manuscript (1998 November). [10] Girault, M., Ton, P. and Vallee, B.: Computation of Approximate Lth Roots Modulo n and Application to Cryptography, Proc. of Crypto'88, LNCS 403, SpringerVerlag, pp (1990). [11] S. Goldwasser, S. Micali and R. Rivest, \A Digital Signature Scheme Secure Against Adaptive ChosenMessage Attacks," SIAM J. on Computing, 17, pp.281{ 308, [12] IEEE P1363 Draft, (1999). [13] Okamoto, T.: A Fast Signature Scheme Based on Congruential Polynomial Operations, IEEE Trans. on Inform. Theory, IT36, 1, pp (1990). [14] Okamoto, T. and Shiraishi, A.: A Fast Signature Scheme Based on Quadratic Inequalities, Proc. of the ACM Symposium on Security and Privacy, ACM Press (1985). [15] Peralta, R.: Bleichenbacher's improvement for factoring numbers of the form N = PQ 2 (private communication) (1997). [16] Peralta, R. and Okamoto, E.: Faster Factoring of Integers of a Special Form, IEICE Trans. Fundamentals, E79A, 4, pp (1996). [17] Pollard, J.L.: Manuscript (1997). [18] Vallee, B., Girault, M. and Ton, P.: How to Guess Lth Roots Modulo n by Reducing Lattice Bases, Proc. of Conference of ISSAC88 and AAECC6 (1988). 7
8 Appendix Security of Ecient Digital Signature Scheme TSHESIGN Eiichiro Fujisaki Tatsuaki Okamoto 1 Introduction In this manuscript, we will prove the security of the ecient digital signature scheme TSHESIGN. 2 Denitions Denition 2.1 Let A be aprobabilistic algorithm and let A(x 1 ;:::;x n ; r) be the result of A on input (x 1 ;:::;x n ) and coins r. We dene by y A(x 1 ;:::;x n ) the experiment of picking r at random and letting y be A(x 1 ;:::;x n ; r). IfS is a nite set, let y be the operation of picking y at random and uniformly from nite set S. " denotes the null symbol and, for list, " denote the operation of letting list be empty. Moreover, jj denotes the concatenation operator, j1jdenotes the size of bit length, i.e., jyj := blog 2 yc +1. For nbit string x, [x] k and [x] k denote the most and least signicant k bits of x, respectively (k n). Denition 2.2 [Random Oracle Model] We dene by the set of all maps from the set f0; 1g 3 of nite strings to the set f0; 1g 1 of innite strings. Let H be a map from a set of an appropriate nite length (say : f0; 1g a )toasetofanappropriate nite length (say f0; 1g b ). H means that map H is chosen from at random and uniformly, with restricting the domain to f0; 1g a and the range to the rst b bits of output. Denition 2.3 [Digital Signature Scheme] A digital signature scheme is dened by a triple of algorithms (Gen; Sig; V er) such that Key generation algorithm Gen is a probabilistic polynomialtime algorithm which on input 1 k (k 2 N) outputs a pair (pk; sk), called public key and secret key. R S 8
9 Signing algorithm Sig is a probabilistic polynomialtime algorithm which on input sk Gen(1 k ) and a message m 2f0; 1g 3 produces a string, s 2f0; 1g 3,called the signatureofm. Verication algorithm Veris a probabilistic polynomialtime algorithm which on input pk Gen(1 k ) and a pair of a message and a signature, (m; s) returns 0 (i.e. invalid) or 1 (i.e. valid) to indicate whether or not the signature is valid. Here we insist that, for any (m; s) where s Sig sk (m), werequire Ver pk (m; s) = 1, otherwise Ver pk (m; s) =0with overwhelming probability. Denition 2.4 [Forging Algorithm F ] Let 5:=(Gen; Sig; V er) be a digital signature scheme. Let A be an algorithm which on input pk obtained by running the generator, Gen, hasaccess to random hash function H and signing oracle Sig sk and eventually outputs some string. We say that adversary A is a (t; q h ;q s ;(k))forger for 5(1 k ) if AdvF 5 (k) := Pr[H ; (pk; sk) Gen(1k ):F H;Sig sk (pk)! (m; s)] (k); where Ver pk (m; s) =1and m is not included in the list of queries that F asks to Sig sk (1), and, moreover, F completes within running time t, asking at most q h queries to H(1) and at most q s queries to Sig sk (1). 3 TSHESIGN: ESIGN with Trisection Size Hash This section introduces the digital signature scheme TSHESIGN. 3.1 Trisection Size Hash function h Although it is necessary that the hash function used in our proposed signature scheme be ideal (or random oracle) to prove the security of the signature scheme, an arbitrary hash function (e.g., MD5 and SHA1) is available for practical use. When the signature scheme is dened over Z=nZ, where jnj = 3k, the underlying hash function h(1) isdenedash : f0; 1g 3!f0; 1g k01. (That is, the size of the image of hash h is around one third of the size of jnj.) 9
10 3.2 Key Generation algorithm Gen Key generation algorithm Gen outputs, given security parameter 1 k (k 2 N), (pk; sk) where pk = fn; eg and sk = fn; e; p; qg. The parameters, n; e; p and q, satisfy the following conditions: p; q (p 6= q) are prime numbers with kbit length, i.e., k = jpj = jqj. n := p 2 q with 3kbit length and e> Signature Generation Sig sk (m) The signature s of message m is computed by the signature algorithm Sig sk (1) as follows: Step 1 Pick r at random and uniformly from (Z=pqZ)npZ := fr 2 Z=pqZj gcd(r; p) =1g. Step 2 Set z (0jjh(m)jj0 2k ) and (z 0 r e )modn. Step 3 Set (w 0 ;w 1 )such that If w 1 2 2k01, then go back to Step 1. w Step 4 Set t 0 er mod p, ands (r + tpq) modn. e01 Step 5 Output s. w 0 = d e; pq (1) w 1 = w 0 1 pq 0 : (2) 3.4 Signature Verication Ver pk (m; s) For a pair of message m and signature s, the verication algorithm Ver pk (m; s) outputs valid (or1)if [s e mod n] k+1 == 0jjh(m)jj0; (3) otherwise invalid (or 0). Here we dene by Ver(h(m);pk) the set of valid signatures, namely, given (h(m);pk), signatures that satisfy Equation (3). 10
11 Remark: In practice, the verication equation can be replaced by [s e mod n] k == 0jjh(m): (4) The formal proof in the next section still works with minor modication of the approximate eth root assumption. 4 Security In this section, we examine the security of TSHESIGN. We rst introduce a problem that is considered to be dicult to solve and then prove that breaking our scheme is as dicult as solving this problem. We call the underlying problem the approximate eth root problem (AERP). Roughly speaking, AERP is the problem, for given (y; e), of nding x such that x e mod n is approximately equivalent toy, i.e., x e mod n y. Now let us dene the approximate eth root problem (AERP) and the breaking algorithm. Denition 4.1 [Approximate eth root problem (AERP)] Let Gen be the generator of the TSHESIGN algorithm. Approximate eth root problem (AERP) is, for given pk := fn; eg 0jjyjj0 ==[x e mod n] k+1. Gen(1 k ) and y R f0; 1g k01, to nd x 2 (Z=nZ)npZ such that Denition 4.2 [AERPBreaker B] We say that adversary B is a (t; (k))aerpbreaker if Adv AERP B (k) := Pr[(pk; sk) Gen(1 k ); y R f0; 1g k01 : B(pk; y)! x] (k); where 0jjyjj0 ==[x e mod n] k and B completes within running time t. The following theorem is the main result of this manuscript. This theorem implies that breaking AERP is as dicult as forging TSHESIGN (existentially forging under adaptive chosen message attacks). 11
12 Theorem 4.3 If there exists a (t(k);q h (k);q s (k);(k))forger F for TSHESIGN, then there exists a (t 0 (k); 0 (k))aerpbreaker B such that Proof: t 0 (k) = t(k)+4(q h + q s ) 1 (k 3 ); and 0 (k) = (1=q h ) 1 (k): Let 5 := (Gen; Sig; V er) be a triple of the TSHESIGN algorithm and F be a (t; q h ;q s ;)forging algorithm for TSHESIGN. We will present an AERPBreaker B that includes F as an oracle. The aim of B is, given y, to nd an approximate root x such that 0jjyjj0 ==[x e mod n] k+1. Recall that the advantage of ESIGNbreaker B is dened by Adv AERP B (k) :=Pr[(pk; sk) Gen(1 k ); y R f0; 1g k01 : B(pk; y)! x]: Likewise recall that the advantage of forger F for TSHESIGN is dened by Adv 5 F (k) :=Pr[H ; (pk; sk) Gen(1k ):F H;Sig sk (pk)! (m; s); where s 2 Ver(H(m);pk)]; where Ver(H(m);pk) is dened as above in Sec.3.4 and m is not included in the list of queries that F asks to Sig sk (1). Since forging algorithm F will make two kinds of queries, hash oracle queries and signing oracle queries, B must answer these queries by itself. Below we describe the specication of AERPbreaker B: AERPBreaker: B(pk; y) set " (empty) and x " (null); set l R f1;:::;q h g, i 0 and j 0; run F(pk); do while F does not ask query Q to H(1) nor ask to Sig sk (1) if F asks query Q to H(1) i ++;j + +; (increment i; j); if j == l else if Q 62 set Q l Q; put (Q l ;";y) in the list and answer F with y; set Q i Q; 12
13 if Q i = Q l abort F and break; do while 0jj 3 jj0 ==[x e i mod n]k+1, x i R (Z=nZ)npZ; set (0jjy i jj0) [x e i mod n]k+1 ; put (Q i ;x i ;y i ) in the list and answer F with y i ; else (Q 2 ) answer F with the corresponding y 0 2 ; else if F asks query Q to Sig sk (1) i + + (increment i); else if Q 62 set Q i Q; End. do while 0jj 3 jj0 ==[x e i x i R (Z=nZ)npZ; mod n]k+1 set (0jjy i jj0) [x e i mod n]k+1 ; put (Q i ;x i ;y i ) in the list and answer F with x i ; else (Q 2 ) answer F with the corresponding x 0 2 ; if F outputs (Q l ;x l ) set x x l ; return x Here '3' denotes an (k 01)bit arbitrary string and denotes a list that records queries of F and the corresponding signatures and hash values that B makes. B starts with list empty and then records (Q i ;x i ;y i )in by following the specication until aborting F or F ends asking queries. Hereafter we explain the strategy of this specication, which is similar to that of FDHRSA as presented by Bellare and Rogawayin[4]. First B picks up l uniformly from f1;:::;q H g. B sets counters, i; j, asi 0and j 0. Counter i is utilized for counting the total number of the queries that F asks to random oracle H and signing oracle Sig, while counter j is utilized for counting the number of F 's asking to H. 13
14 If F asks query Q to random oracle H(1), B increases i; j and then works as follows: { If the query is the lth one, B sets Q l Q, puts (Q l ;";y) in list, and answers F with y, where y is the instance input to B, { Else if the query has not been recorded in list, B sets Q i Q, and then after executing Experiment I= [repeatedly picks up x i from (Z=nZ)npZ until 0jj 3 jj0 ==[x e i mod n]k+1, where '3' denotes an (k 0 1)bit arbitrary string], makes an entry of (Q i ;x i ;y i )in and answers F with y i as the hash value of Q, where y i is dened by 0jjy i jj0 [x e i mod n]k+1. { Otherwise (i.e., Q is already in list ), B answers F with the corresponding y 0 2. Similarly, iff asks query Q to signing oracle Sig sk (1), B increases i and then works as follows: { If Q i = Q l, B aborts F and outputs " (This means B failed to solve AERP), { Else if Q 62, B sets Q i Q, and then after executing Experiment I makes an entry of (Q i ;x i ;y i )in and answers F with x i as the signing value of Q, where y i is dened by 0jjy i jj0 [x e i mod n]k+1. { Otherwise (i.e., Q is already in list ), B answers F with the corresponding x 0 2. Here we state that the distribution of (x i ;y i ) from Experiment I is identical to that of (s; H(m)) of signing algorithm Sig sk (1) in the random oracle model. To prove this, we state the following two lemmas. To prove Lemma 4.5 is sucient to prove this statement. Lemma 4.4 For given h(m) 2f0; 1g k01 and pk holds: Gen(1 k ), the following equation #Ver(h(m);pk)=#fr 2 (Z=pqZ)npZj0 (0 mod pq)+pq < 2 2k01 g; (5) where := ((0jjh(m)jj0 2k ) 0 r e )modn. 14
15 Sketch of Proof: Recall that Ver(h(m);pk):=fs 2 (Z=nZ)npZj[s e mod n] k+1 == 0jjh(m)jj0g. Represent s 2 Ver(h(m);pk)bys = r + tpq where r 2 (Z=pqZ)npZ, andt 2 Z=pZ. Itis easy to check that (r; t) is the unique representation of s. Let Setr(h(m)) := fr 2 (Z=pqZ)npZj0 (0 mod pq) +pq < 2 2k01 g where := ((0jjh(m)jj0 2k ) 0 r e )modn. This makes it possible to make bijective map Ver(h(m);pk)! Setr(h(m)) by r := s mod pq and t := (0 mod pq) +pq where := ((0jjh(m)jj0 2k ) 0 r e )modn. Therefore, #Ver(h(m);pk)=#Setr(h(m)). { Lemma 4.5 For given s 0 2 Ver(h(m);pk), the following equation holds: Pr[s Sigsk(m) h :s == s 0 ]= Pr [s == s 0 js 2 Ver(h(m);pk)] = s R(Z=nZ)npZ where Pr s R(Z=nZ)npZ [s == s 0js 2 Ver(h(m);pk)] denotes the conditional probability of Event [s R (Z=nZ)npZ : s == s 0 ] given Event [s R (Z=nZ)npZ : s 2 Ver(h(m);pk)]. Sketch of Proof: >From Lemma 4.4, the proof of this lemma is straightforward. { >From Lemma 4.5, we found that B can answer F by itself unless F asks Q l as a signing query. Eventually, when F outputs (Q; s), Q is considered to be Q j for some j. In the case of j = l, B succeeds to break AERP because [s e mod n] k+1 == y, and the success probability 0 (k) isatleast(1=q h ) 1 (k). { 1 #Ver(h(m);pk) ; (6) 15
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