Patent Application: US-201214110448-A

Abstract:
methods and apparatus for generating a private - public key pair , for encrypting a message for transmission through an unsecure communication medium , and for decrypting the message are disclosed . the methods are based on the well - known mceliece cryptosystem or on its niederreiter variant . more general transformation matrices q are used in place of permutation matrices , possibly together with an appropriate selection of the intentional error vectors . the transformation matrices q are non - singular n × n matrices having the form q = r + t , where the matrix r is a rank - z matrix and the matrix t is some other matrix rendering q non - singular . the new q matrices , though at least potentially being dense , have a limited propagation effect on the intentional error vectors for the authorized receiver . the use of this kind of matrices allows to better disguise the private key into the public one , without yielding any further error propagation effect . based on this family of q matrices , the presently proposed cryptosystem enables the use of different families of codes than goppa codes , such as rs codes , by ensuring increased public key security .

Description:
a preferred embodiment of the present invention is explained in the following with reference to fig2 and 3 . bob operates a key generating apparatus 10 . this apparatus comprises a private - key generating module 11 to generate bob &# 39 ; s private ( secret ) key in the form of three matrices . one of these matrices is a k × n systematic generator matrix g of a linear block code ( code length n , code dimension k and capability to correct t errors ) defined over gf ( p ). the other two matrices are a k × k non - singular scrambling matrix s and an n × n non - singular transformation matrix q . a public key is then generated in a public - key generating module 12 . the public key comprises two parts : a k × n matrix of the form and optionally includes a set of error constraints as will be explained below . in general , in contrast to the original mceliece cryptosystem , the public code is not permutation - equivalent to the private code due to a specific choice of q , as will also be explained below . bob publishes the public key in a suitable format with a publication module 13 . this module may comprise an e - mail sending device to communicate the public key to selected users , a public directory accessible through the internet or through any other network by protocols such as http or ftp , or any other means to distribute the public key to potential transmitters of encrypted information . alice encrypts a message represented by a cleartext vector u by using an encryption apparatus 20 . to this end , a key receiving module 21 obtains bob &# 39 ; s public key . an error - generation module 22 produces an error vector e , which is a 1 × n vector of t intentional errors that may be subject to the error constraints optionally included in the public key , as explained further below . a ciphertext - generating module 23 then applies the following encryption map : where x is a ciphertext vector corresponding to the cleartext u . the ciphertext is then transmitted by a transmitter 24 through an unsecure communication medium 30 , as , for example , through the internet or through a wireless communication link , or stored onto a storage medium , as , for example , a hard disk or flash storage device . the communication medium 30 may be the same as the one through which the public key was transmitted , or it may be a different medium . bob receives and decrypts the message with a decryption apparatus 40 . the decryption apparatus 40 comprises a receiver 41 for receiving the ciphertext vector from the communication medium 30 . a cleartext deciphering module 42 reconstructs a cleartext vector u from the ciphertext vector x by sending the ciphertext vector to a first multiplication submodule 421 , which inverts the transformation as follows : thus obtaining a product vector x ′ representing a codeword of the secret code affected by the error vector e · q . a decoding submodule 422 will then decode the product vector , correcting all the errors , and get u · s − 1 , due to the systematic form of g . error correction of this type is well known in the prior art and will depend on the linear block code which is actually employed . a second multiplication submodule 423 then multiplies u · s − 1 by s to obtain the reconstructed cleartext vector u . this vector is then outputted in a suitable form by an output module 43 . in general , the use of a transformation matrix q in place of a permutation matrix causes an error propagation effect . however , if q is chosen within a given class of matrices , this effect can be controlled or even eliminated . this can be achieved by giving q the following form : with r a ( preferably dense ) n × n matrix having rank z & lt ; n and t a ( preferably sparse ) n × n matrix such that q is non - singular . the optional error constraints minimizing the error propagation can then be expressed as e · r = 0 . in this case , when implementing decryption , bob experiences a residual error vector e · q = e · t . a possible form for r is the following : where a and b are two random z × n matrices and t denotes transposition . in this case , the intentional error vectors should be restricted to only those vectors satisfying the condition a · e t = 0 , and the matrix a should be disclosed as the second part of the public key . in order to increase the system security , the error constraints can be omitted from the public key . in this case , no selection of the intentional error vectors is performed and the product e · r can assume whatever value . in order to successfully complete decryption , the value of e · r can be found by the legitimate receiver by first guessing it and then performing a set of iterated decryption attempts , until one of them succeeds . as far as matrix t is concerned , a possible form for it is the following : t = ∑ i = 1 m ⁢ π i , ( eq . ⁢ 6 ) where each π i is a generalized permutation matrix , having only one non - zero element in each row and column , whose value is randomly selected among the p − 1 non - zero elements of gf ( p ). this gives a matrix t with , at most , m non - zero elements in each row and column . the usage of t as in ( eq . 6 ) leads to an increase in the number of intentional errors by ( at most ) a factor m . such controlled error propagation effect can be compensated by bob by using codes with a high error correction capability , like ldpc codes [ 10 ]. in order to reduce the amount of storage needed for representing the public key , linear block codes having quasi - cyclic nature can be used in the present invention . in such case , n and k are multiple of an integer q ( n = n 0 × q , k = k 0 × q ) and matrices s , g , r , t are formed by circulant sub - matrices with size q × q . when qc codes are adopted , the design of matrices r and t can be performed block - wise , still aiming at obtaining a full - rank matrix q = r + t . suitable circulant matrices and circulant permutation matrices can be designed to form the blocks of matrices r and t , so ensuring the qc nature of matrix q , that , together with that of s and g , suffices to obtain a public key g ′ in qc form . in such case , g ′ can be represented by only storing the first row ( or column ) of each one of its circulant blocks ; so , the key length is reduced with respect to the use of non - qc codes and it increases linearly with the code length . when using qc codes , the optional set of constraints on the intentional error vectors can be imposed block - wise , in such a way that each q - bit sub - vector of an error vector must verify one or more error constraints . in a simplified embodiment , obtained by using r and t matrices that are special cases of equations ( 5 ) and ( 6 ), the matrix q takes the following form : r = a t · b =[ a 1 a 2 . . . a n ] t ·[ b 1 b 2 . . . b n ] ( eq . 8 ) in the choice of q 1 it is preferred to avoid some special cases which could allow an attacker to obtain a code that is permutation - equivalent to the secret one , thus bringing security back to that of the classical mceliece cryptosystem . suppose the j - th element of b is zero and that p 1 has a symbol 1 at position ( i , j ). in this case , the j - th column of q 1 is completely null , except for its element at row i . since q 1 − 1 ={ circumflex over ( q )} 1 /| q 1 |, where { circumflex over ( q )} 1 is the adjoint matrix and | q 1 | is the determinant of q 1 , it follows from the definition of { circumflex over ( q )} 1 that the i - th column of q 1 − 1 is completely null , except for its element at row j ( that is not necessarily equal to 1 , except for the binary case ). so , the i - th column of q 1 − 1 has the effect of a column permutation ( apart from multiplication by a constant ), like in the original mceliece cryptosystem . in order to avoid such a case , it should preferably be imposed that all the elements of b are non - zero . such condition could not be sufficient in the binary case , for which it translates into having b that is the all - one vector . in this case , if a is an all - one vector too , i . e ., r = 1 . it is known that a binary matrix q 1 = 1 + p 1 is invertible only when it has even size and , in such case , it can be shown that its inverse has the form q 1 − 1 = 1 + p 1 t , due to the property of permutation matrices ( as orthogonal matrices ) to have their inverse coincident with the transpose . so , q 1 − 1 has the same form of q 1 and disclosing g ′= s − · g · q 1 − 1 might imply disclosing a generator matrix of a permuted version of the secret code or its complementary ( depending on the parity of its row weight ). therefore , the form q 1 = 1 + p 1 might reduce the security to that of the permutation used in the original mceliece cryptosystem . restricting to the binary case with q in the special form of eq . ( 7 ), even adopting a vector a different from the all - one vector would not suffice to guarantee increased security with respect to the classical mceliece cryptosystem . since q 1 − 1 ={ circumflex over ( q )} 1 /| q 1 |, it is easy to verify that a weight - 1 row in q 1 produces a weight - 1 row in q 1 − 1 and a weight -( n − 1 ) row in q 1 produces a weight -( n − 1 ) row in q 1 − 1 . it follows that q 1 − 1 contains couples of columns having hamming distance 2 . since their sum is a weight - 2 vector , the sum of the corresponding columns of the public matrix results in the sum of two columns of s − 1 g . starting from this fact , an attacker could try to solve a system of linear equations with the aim of obtaining a permutation - equivalent representation of the secret code , at least for the existing distance - 2 column pairs . so , in the special case of the binary field , it is preferable to avoid the choice of q as in ( eq . 7 ). a better choice of q is obtained by using an r matrix with rank z & gt ; 1 and by adding more than one permutation matrix to it ( i . e . m & gt ; 1 ). cryptosystems based on full decoding [ 15 ], [ 16 ] aim at increasing the hardness of the basic problem and , for this purpose , a transformed set of intentional error vectors must be adopted , thus requiring to partially disclose the secret transformation . in contrast , the system according to the present invention still bases its robustness on the same problem as the original mceliece cryptosystem ( that is , normal decoding ); so , the error vectors are the same used in the original mceliece cryptosystem , apart from an optional selection of them , while no transformation of the error vectors is needed . for this reason , the information leakage on the secret transformation that is needed in the cryptosystem according to the present invention is considerably lower with respect to what happens in full decoding cryptosystems . in the first version of full decoding cryptosystem [ 15 , sec . 8 . 3 ], p rows of the secret transformation matrix were made public . in the system according to the present invention , the secret transformation matrix may be chosen as in eq . ( 4 ), with r as in eq . ( 5 ), and only a can be disclosed or , even more notably , no error constraint can be needed when the legitimate receiver exploits a set of iterated decryption attempts . the cryptosystem object of the present invention can be alternatively implemented through the niederreiter variant of the mceliece cryptosystem [ 4 ], that is able to exploit shorter public keys with respect to the classical mceliece implementation . when the niederreiter implementation is preferred , the first part of the public key is an r × n parity - check matrix h ′ of the form where the matrix s is a non - singular r × r matrix over gf ( p ), the matrix h is an r × n parity - check matrix of the private ( n , k ) linear block code over gf ( p ) and matrix q is a non - singular n × n matrix as in ( eq . 4 ). in this case , the ciphertext x has size r × 1 and is obtained from the cleartext u by mapping u into a 1 × n weight t error vector e and calculating its syndrome x through h ′, that is : vector e may be optionally selected among the weight t error vectors verifying the condition e · r = 0 . when the ciphertext is received by the legitimate receiver , it is multiplied by s , so obtaining : performing syndrome decoding on x ′ ( or on a vector derived from x ′, when no selection of the error vectors is performed ) with the secret linear block code ( represented by h ) allows to obtain the vector t t · e t , that is then multiplied by ( t t ) − 1 , so recovering e . the latter can be hence demapped into its associated cleartext vector u . the security level is commonly measured as the number of elementary operations ( or work factor ) of the most efficient attack . with reference to ref [ 3 ], to reach a security level & gt ; 2 80 , the ( 1632 , 1269 ) goppa code is suggested , resulting in a public - key size of 460647 bits ( that is the lowest possible value for this code , obtained by storing the non - systematic part of h , as in the niederreiter cryptosystem ). with the variant according to the present invention , we can include the rs code with n = 255 , k = 195 , t = 30 , having a security level of 2 84 . 18 ( estimated as the work factor of the attack procedure described in [ 17 ]). the public key size for this code , if we store the 195 × 255 matrix g ′ and a 1 × 255 vector a , both with elements over gf ( 256 ), is ( 195 + 1 )· 255 · 8 = 399840 bits , that is , about 13 % less than ( the minimum size of ) that obtained by the revised mceliece cryptosystem [ 3 ]. by exploiting the niederreiter variant of the system according to the present invention , a systematic h ′ can be used as the public key , and it can be represented by only storing its non - systematic part , that is , 95640 bits . this results in a reduction of about 79 % with respect to the goppa - based solution . the security level of the two systems remains comparable even when the constraint expressed by a is imposed on the intentional error vectors of the modified cryptosystem . in fact , the introduction of each constraint results in a decreased security level of about 2 3 . we consider the goppa code suggested in ref [ 3 ] to achieve a security level ≧ 2 128 which has n = 2960 , k = 2288 , yielding a key length of 1537536 bits . in the system according to the present invention , we can adopt an rs code with the same rate ( 0 . 77 ), defined over gf ( 512 ), having n = 511 , k = 395 , t = 58 . the corresponding size of the public key , including a 1 × 511 vector a , is ( 395 + 1 )· 511 · 9 = 1821204 bits ( that is slightly bigger than the one in the niederreiter system proposed in ref [ 4 ]). by exploiting the niederreiter variant of the system according to the present invention , a systematic h ′ can be used as the public key , and it can be represented by only storing its non - systematic part , that is , 416979 bits . this results in a reduction of about 73 % with respect to the goppa - based solution . furthermore , by using the present solution , the security level grows up to 2155 . 89 ( estimated on the basis of the attack procedure described in ref [ 17 ]). this value remains very high even when we consider the presence of the constraint expressed by a on the intentional error vectors . in some of the above examples , the matrix r was chosen as r = a t · b , and the intentional error vectors were restricted to only those vectors satisfying the condition a · e t = 0 , meaning that the matrix a had to be disclosed as part of the public key . such a choice was made both for the scheme which follows the original mceliece scheme , and for the niederreiter variant . it can be shown that such a setting exposes a subcode of the public code that is permutation - equivalent to a subcode of the private code . in fact , referring to the niederreiter variant , an attacker could consider the subcode generated by the following parity - check matrix : each codeword c in the code defined by h ″ must fulfil the condition a · c t = 0 . this also implies r t · c t = 0 ; so , matrix h ″ selects those codewords c such that s − 1 h · t t · c t = 0 . hence , the effect of the dense r is removed and , when t is a permutation matrix , the subcode defined by h ″ is permutation - equivalent to a subcode of the secret code . in such case , when rs codes are used , an attacker could implement an attack of the type described in ref . [ 6 ]. for practical choices of the system parameters , it seems that the subcode resulting by the inclusion of matrix a in h ″ is always weak against such attack . based on these premises , some possible countermeasures can be designed , that are able to make the system more robust against this kind of attacks , by trading an increased security level for a higher complexity of the decryption stage . the rationale of these variants is : 1 . breaking the straight connection between r and the public representation of the error constraints . 2 . exploiting the secret nature of the vector b , which is uninfluent when r = a t · b and a · e t = o . by referring to the case in which a is a vector ( but the method can also be used when a is a matrix ), a first countermeasure consists in splitting the constraint vector a into two ( or more ) subvectors . one may consider the case of splitting into two subvectors by choosing the constraint vector a is still made public , but the condition on the error vectors is changed into ( this may be generalized to a · e t = γ , wherein γ is an element of gf ( p )). alternatively , the constraint vector a can be kept secret . in this case , the value of a · e t must be guessed by the legitimated receiver and decryption must be tried with all the possible values of a · e t , until success . in order to break its connection with a , the matrix r is modified as follows : r = [ a 1 a 2 ] t · [ b 1 - b ] , ( eq . ⁢ 13 ) with r in the form ( eq . 13 ), the value of r t · e t can easily be obtained starting from the value assumed by a i · e t = α . in this case , a 2 · e t = 1 − α and this may also be generalized to the case in which , in the place of 1 − b , a further secret vector d is introduced , that must be randomly chosen . by knowing ( or guessing , as can be seen in the following ) the value of a , the authorized receiver can exploit the knowledge of b ( and of d , if used in the place of 1 − b ) to calculate r t · e t and decrypt the received ciphertext . concerning the attack , with r in the form ( eq . 13 ), the matrix h ″ as in ( eq . 12 ) no longer defines a subcode which is permutation - equivalent to a subcode of the secret code , so the original version of the attack is prevented . however , eve could try to sum , at random , two rows of h ″, hoping that one of them corresponds to a copy of the vector a 1 in r and the other to a copy of the vector a 2 . in this case , the sum of the two rows would still contain the vector a , so being suitable for the selection of the weak subcode . however , this approach would produce , on average , only r / 2 =( n − k )/ 2 rows containing vector a , while the others would contain 2a 1 or 2a 2 . furthermore , the sum of couples of rows would also imply to sum couples of rows of t t ; so , their ( very low ) weight would be doubled with a very high probability . for these reasons , it seems not possible to devise a variant of the subcode attack able to be successful on this modified setting . in order to describe how the decryption stage must be modified according to this new variant , a distinction will be made between the binary and non - binary cases in the following . in the binary case , according to ( eq . 14 ), r t · e t can only assume two values : { α = 1 ⇒ r t · e t = b t , α = 0 ⇒ r t · e t = ( 1 + b ) t . suppose that bob first computes x ′= s · x = h · q t · e t = h · r t · e t + h · t t · e t and then , by guessing that α = 1 , he computes x 1 ″= x ′ h · b t . if it is actually α = 1 , then r t · e t = b t and x 1 ″= h · t t · e t . so , bob can perform successful syndrome decoding on x 1 ″ to recover e . in order to verify that e is actually that transmitted by alice , bob can check its weight and also verify that a 1 · e t = α = 1 . if instead α = 0 , then r t · e t =( 1 + b ) t and x 1 ″= h · 1 t + h · t t · e t . if the code is self - complementary , then h · 1 t = 0 and bob can proceed exactly in the same way as for α = 1 . otherwise , there are two possibilities : i ) x 1 ″ does not correspond to any correctable syndrome , so decoding fails or ii ) x 1 ″ corresponds to a correctable syndrome , so decoding results in an error vector e ′. in the latter case , bob can check the weight of e ′ and also verify whether a 1 · e ′ t = α = 1 . this suffices to detect that e ′ is not that transmitted by alice . in such case , bob knows that α = 1 and he must restart from x 0 ′= x ′+ h ·( 1 + b ) t . by this procedure , the probability of an undetected error , that is , e ′= e , is comparable to the probability of undetected error that is typical for the secret code . for the kind of codes considered in the cryptosystem , the probability of such an event is extremely low . so , for the binary case , bob is always able to distinguish between the two cases α = 0 and α = 1 , and to correctly decrypt the secret message . by using the proposed method , based on guessing the value of α , bob needs , on average , 1 . 5 decryption attempts on each received message . so , this variant of the cryptosystem , for the binary case , has decryption complexity increased , at most , by a factor 2 0 . 6 . instead , if the code is self - complementary , bob does not need to distinguish between α = 0 and α = 1 ; so , the decryption complexity remains unchanged . for the case of gf ( p ), there are obviously p possible values of α . so , bob needs to make up to p guesses on the value of α . first , bob computes x ′= s · x = h ·( r + t ) t · e t . by using ( eq . 14 ), this results in then , bob makes his first guess by supposing α = α b and calculates x ′ α b = x ′− h ·[ α b b t +( 1 − α b )( 1 − b ) t ] ( eq . 16 ) by replacing ( eq . 15 ) in ( eq . 16 ) and by some simple arithmetics , one obtains x ″ α b = h · t t · e t + 2 ( α − α b ) h · b t −( α − α b ) h · 1 t . ( eq . 17 ) if α b = α , then x ″ α b = h · t t · e t ; so , bob can recover e through syndrome decoding , check its weight and verify that a 1 · e t = α b . otherwise , applying syndrome decoding on x ″ α b expressed by ( eq . 17 ) always results in a decoder failure or in obtaining e ′≠ e , both in the case of self - complementary codes or not . the latter case can be detected by bob by checking the weight of e ′ and verifying that a 1 · e ′ t ≠ α b . so , the modified version can be used also in the non - binary case . obviously , when working over gf ( p ), the average number of decryption attempts needed by bob becomes ( p + 1 )/ 2 , that also gives a measure of the increase in the decryption complexity . actually , some steps of the decryption procedure do not need to be repeated ; so , an increase in the decryption complexity by a factor of ( p + 1 )/ 2 is a worst case estimation . starting from the constraint vector splitting approach described in the previous section , an even simpler countermeasure can be obtained , consisting in maintaining the original setting of the proposed cryptosystem ( i . e . with r = a t · b ), but hiding the constraint vector a . this obviously eliminates the need of selecting the intentional error vectors . this second countermeasure is illustrated for the case of a and b being two 1 × n vectors , but can be generalized to the use of a and b in the form of two z × n matrices , with z & gt ; 1 . suppose that a is private and that the error vector e generated by alice is such that a · e t = γ , with γεgf ( p ). it follows that in this case , bob can guess that the value of γ is γ b and compute so , if γ b = γ , bob obtains x ″= h · t t · e t . in such case , he can recover e through syndrome decoding , check its weight and verify that a · e t = γ b . otherwise , supposing that b is not a valid codeword , syndrome decoding fails or returns an error vector e ′≠ e . this latter case can be identified by bob by checking the weight of e ′ and the value of a · e ′ t . the probability of finding e ′= e , for γ ≠ γ b , is extremely low . also in this case , the complexity of the decryption stage is increased , on average , by a factor ≦( p + 1 )/ 2 . whereas both the first and the second countermeasure have been described in the context of the niederreiter variant of the proposed cryptosystem , these countermeasures are equally applicable to encryption and decryption following the scheme , described above , which is similar to the original mceliece scheme . r . j . mceliece , “ a public - key cryptosystem based on algebraic coding theory .” dsn progress report , pp . 114 - 116 , 1978 . e . berlekamp , r . mceliece , and h . van tilborg , “ on the inherent intractability of certain coding problems ,” ieee trans . inform . theory , vol . 24 , no . 3 , pp . 384 - 386 , may 1978 . d . j . bernstein , t . lange , and c . peters , “ attacking and defending the mceliece cryptosystem ,” in post - quantum cryptography , ser . lecture notes in computer science . springer berlin / heidelberg , 2008 , vol . 5299 / 2008 , pp . 31 - 46 . h . niederreiter , “ knapsack - type cryptosystems and algebraic coding theory ,” probl . contr . and inform . theory , vol . 15 , pp . 159 - 166 , 1986 . r . overbeck , “ structural attacks for public key cryptosystems based on gabidulin codes ,” journal of cryptology , vol . 21 , no . 2 , pp . 280 - 301 , 2008 . c . wieschebrink , “ cryptanalysis of the niederreiter public key scheme based on grs subcodes ,” in post - quantum cryptography : pqcrypto 2010 , ser . lncs , n . sendrier , ed . springer , 2010 , vol . 6061 , pp . 61 - 72 . v . g . umana and g . leander , “ practical key recovery attacks on two mceliece variants ,” in proc . 2nd int . conf . on symbolic computation and cryptography , c . cid and j .- c . faugere , eds ., egham , uk , june 2010 , pp . 27 - 44 . c . monico , j . rosenthal , and a . shokrollahi , “ using low density parity check codes in the mceliece cryptosystem ,” in proc . ieee international symposium on information theory ( isit 2000 ), sorrento , italy , june 2000 , p . 215 . a . otmani , j . p . tillich , and l . dallot , “ cryptanalysis of two mceliece cryptosystems based on quasi - cyclic codes ,” in proc . first international conference on symbolic computation and cryptography ( scc 2008 ), beijing , china , april 2008 . m . baldi , m . bodrato , and f . chiaraluce , “ a new analysis of the mceliece cryptosystem based on qc - ldpc codes ,” in security and cryptography for networks , ser . lecture notes in computer science . springer berlin / heidelberg , 2008 , vol . 5229 , pp . 246 - 262 . e . m . gabidulin , a . v . paramonov , and o . v . tretjakov , “ ideals over a noncommutative ring and their application in cryptography ,” d . w . davies , ed ., advances in cryptology — eurocrypt 91 , lecture notes in computer science , vol . 547 , pp . 482 - 489 , 1991 . t . p . berger and p . loidreau , “ how to mask the structure of codes for a cryptographic use ,” designs , codes and cryptography , vol . 35 , pp . 63 - 79 , 2005 . a . ourivski and e . gabidulin , “ column scrambler for the gpt cryptosystem ,” discrete applied mathematics , vol . 128 , pp . 207 - 221 , 2003 . h . rashwan , e . m . gabidulin , and b . honary , “ security of the gpt cryptosystem and its applications to cryptography .” security comm . networks , 2010 . g . kabatiansky , e . krouk , and s . semenov , error correcting coding and security for data networks : analysis of the superchannel concept . john wiley & amp ; 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