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BY 4.0 license Open Access Published by De Gruyter November 17, 2020

Algebraic approaches for solving isogeny problems of prime power degrees

Yasushi Takahashi EMAIL logo , Momonari Kudo , Ryoya Fukasaku , Yasuhiko Ikematsu , Masaya Yasuda and Kazuhiro Yokoyama


Recently, supersingular isogeny cryptosystems have received attention as a candidate of post-quantum cryptography (PQC). Their security relies on the hardness of solving isogeny problems over supersingular elliptic curves. The meet-in-the-middle approach seems the most practical to solve isogeny problems with classical computers. In this paper, we propose two algebraic approaches for isogeny problems of prime power degrees. Our strategy is to reduce isogeny problems to a system of algebraic equations, and to solve it by Gröbner basis computation. The first one uses modular polynomials, and the second one uses kernel polynomials of isogenies. We report running times for solving isogeny problems of 3-power degrees on supersingular elliptic curves over 𝔽p2 with 503-bit prime p, extracted from the NIST PQC candidate SIKE. Our experiments show that our first approach is faster than the meet-in-the-middle approach for isogeny degrees up to 310.

1 Introduction

Since Koblitz [25] and Miller [27] independently proposed elliptic curve cryptography (ECC) in 1985, elliptic curves have been used in cryptography. Since 2000, pairings over elliptic curves have been an indispensable tool to construct cryptographic protocols and functional encryption schemes. Since 2006, isogenies between elliptic curves have began to be used in [9, 29, 34] for constructing several cryptosystems and hash functions. In particular, supersingular isogeny graphs were first proposed in [9] for security, which introduced the supersingular isogeny graph path-finding problem as a hard problem in cryptography. Actually, there exists a subexponential quantum algorithm [10] for computing an isogeny between ordinary elliptic curves since it is reduced to solving the abelian hidden shift problem. (Recently, the key exchange protocol based on group action [8] is regarded as a credible post-quantum system despite of the existence of a sub-exponential attack.) In contrast, the supersingular case is non-abelian and seems to be a promising candidate for PQC. (cf., The security of ECC and pairing-based cryptography relies on the discrete logarithm problem over elliptic curves, which could be solved in polynomial-time by Shor’s algorithm [31] using quantum computers.) In 2011, Jao and De Feo [22] introduced the key exchange protocol using using supersingular isogenies as a post-quantum candidate, based on pseudo-random walks in supersingular isogeny graphs. (See [11] for the connection between the hard problems in [22] and the path-finding problem in [9].) Other cryptographic functions were subsequently developed in [14, 23, 35]. In 2017, Jao et al. [21] submitted algorithms of supersingular isogeny key encapsulation, called SIKE, for the NIST PQC standardization process. It remains as a second-round candidate [28].

The template for the security of isogeny cryptosystems is the general isogeny problem [20]; Given two elements j,ȷ˜ of a finite field 𝔽q, to find an isogeny ϕ:EE˜ if exists, such that j(E) = j and j(E˜)=ȷ˜ where j(E) denotes the j-invariant of an elliptic curve E. In the supersingular case, it is sufficient to take 𝔽p2 as the base field for a prime p since every supersingular curve over Fp is isomorphic to one defined over 𝔽p2 [33]. A variant of this problem is when the degree of ϕ is known, and it arises from the cryptanalysis of the hash function of [9], which requires computing isogenies of degree 0e for some small 0 and large e. The most cryptographically-interesting variant is the supersingular isogeny Diffie-Hellman (SIDH) problem [14], assuring the security of [21, 22], in which supersingular elliptic curves over 𝔽p2 are chosen for a special prime p=1e12e2f±1 and a lot of auxiliary information are given (see also [19]). The basic algorithm to solve the general isogeny problem for ordinary curves is due to Galbraith [18]. His procedure is (1) to reduce the problem to the case of elliptic curves E and E˜ whose endomorphism ring is maximal, and then (2) to construct an isogeny between E and E˜ by building isogeny trees. The time complexity of step (1) for E is negligible when the conductor [𝒪K : End(E)] is smooth, where 𝒪K denotes the maximal order of the imaginary quadratic field K=End(E)ZQ with endomorphism ring End(E). For step (2), a sub-exponential quantum algorithm was discovered in [10]. In contrast, the meet-in-the-middle approach in [18] is applicable to the supersingular isogeny graph by building isogeny trees from E to E˜ directly. The best known quantum algorithm is due to Biasse et al. [3] for the supersingular problem, and its time complexity is exponential. For the SIDH problem, a faster quantum algorithm is given by Tani’s claw finding algorithm [36], but its time complexity is still exponential. (See also [24].) In recent, several related computational problems have been discussed in [15] for supersingular elliptic curves, their isogeny graphs, and their endomorphism rings.

Although time complexities of fast quantum algorithms has been discussed actively as described above, we consider practical approaches for solving isogeny problems with classical computers. Here we focus on the supersingular isogeny graph path-finding problem in [9] of prime power degree =0efor small0. The meet-in-the-middle approach is a practical way to solve the isogeny problem with time complexity equal to the square root of . It seems the most practical way at least in the supersingular case. (It is reported in [1] that van Oorschot and Wiener’s parallel golden collision search [38] can outperform the meet-in-the-attack for solving isogeny problems with parallel computation. See also [12] for improvements of the parallel golden collision search.) While the meet-in-the-middle approach is a generic way for solving graph problems, we propose two algebraic approaches for solving the isogeny problem in this paper. (cf., Several algebraic approaches are considered in [26] to attack the SIDH/SIKE protocol.) Our basic strategy is to reduce the isogeny problem to a system of algebraic equations. (It does not depend on the supersingularity.) Specifically, our first approach uses the modular polynomial of prime level 0. We divide a system of equations of modular polynomial into two parts like the meet-in-the-middle approach, and compute their Gröbner bases to efficiently find j-invariants of intermediate curves between given two isogenous elliptic curves EandE˜. In contrast, our second approach takes an intermediate curve E 0 between EandE˜, and consider kernel polynomials F (x) and F˜(x˜) of two isogenies φ:EE0andφ˜:E˜E0. Since the curve E0 is unknown, we regard its Weierstrass coefficients as variables, and also represent kernel polynomials F(x)andF˜(x˜) as certain multivariate polynomials, based on Schoof’s work [30]. Furthermore, by using Vélu’s formulae [6, 39], we represent isogenies φ and φ˜ as rational functions over multivariate polynomial rings, and set up algebraic equations to determine the Weierstrass coefficients of E0. Moreover, we report running times of our algebraic approaches for solving the isogeny problem of 3-power degrees on supersingular elliptic curves over 𝔽p2 with 503-bit prime p, extracted from SIKE-p503 parameters [21], to compare with the meet-in-the-middle approach in performance.

2 Mathematical background

In this section, we review some basic definitions and properties of elliptic curves and isogenies, which shall be required in Section 3 below.

2.1 Elliptic curves over finite fields

An elliptic curve E over a finite field 𝔽q of characteristic p ≥ 5 is defined by the (short) Weierstrass form y2 = x3 + ax + b with a, b ∈ 𝔽q and discriminant Δ(E) = −16(4a3 + 27b2) ≠ 0. Its j-invariant is defined as j(E)=1728(4a)3Δ(E). Two elliptic curves are isomorphic over Fq if and only if they both have the same j-invariant, where Fq denotes the algebraic closure of 𝔽q. Moreover, given an element j0 of 𝔽q, there exists an elliptic curve over 𝔽q with j-invariant equal to j0. The set of 𝔽q-rational points


forms an abelian group, where 𝒪E denotes the infinity point on E. (See [33, Chap. III] for the group law.) The number of 𝔽q-rational points on E, denoted by #E(𝔽q), is finite, and it is represented as #E(𝔽q) = q + 1 − t where t denotes the trace of the qth-power Frobenius map (see [33, Chap. V] for the map). It satisfies |t|2q by Hasse’s theorem. An elliptic curve E over 𝔽q is said supersingular if the characteristic p divides the trace t. Otherwise E is said ordinary. Every supersingular elliptic curve over Fp has its j-invariant defined over 𝔽p2 [33, Thm. 3.1, Chap. V], and hence it is isomorphic (over Fq ) to one defined over 𝔽p2. Furthermore, there are about p12 isomorphism classes of supersingular elliptic curves over Fp.

2.2 Isogenies and Vélu’s formulae

Let EandE˜ be two elliptic curves over a finite field 𝔽q. A morphism ϕ:EE˜ satisfying ϕ(OE)=OE˜ is called an isogeny. Two elliptic curves EandE˜ are called isogenous if there is a non-zero isogeny between them. Tate’s theorem [37] states that EandE˜ are isogenous over 𝔽q if and only if #E(Fqk)=#E˜(Fqk) for any positive integer k. Every non-zero isogeny ϕ:EE˜ induces an injection of function fields ϕ:Fq(E˜)Fq(E) (see [33, Chap. III]). The degree of a non-zero isogeny is defined as deg ϕ=[Fq(E):ϕFq(E˜)]. We say that ϕ is separable if the finite extension Fq(E)/ϕFq(E˜) is separable. A non-zero isogeny is separable if its degree is not divisible by the characteristic of the base field 𝔽q.

A non-zero isogeny ϕ:EE˜ induces a (surjective) group homomorphism from E(Fq) to E˜(Fq) Thm. 4.8, Chap. III], and its kernel is a finite subgroup of E(Fq), denoted by E[ϕ]. It satisfies deg ϕ = #E[ϕ] if ϕ is separable. Conversely, given any finite subgroup SofE(Fq), there are a unique elliptic curve E˜ and a separable isogeny ϕ:EE˜withE[ϕ]=S [33, Prop. 4.12, Chap. III]. The curve E˜ is denoted by the quotient E/S. Vélu [39] showed how to explicitly represent the isogeny ϕ:EE˜=E/S and the Weierstrass equation for E˜. A non-zero isogeny ϕ:EE˜ is normalized if ϕ(ωE˜)=ωE, where ωEandωE˜ denote the invariant differentials of E$and$E˜, respectively (see [33, Chap. III] for the invariant differential of an elliptic curve). Vélu’s formulae give a normalized separable isogeny, and its modified form is shown in [6] as below (the form was given much earlier by Elkies [16]):

Modified Vélu’s formulae in [6]

Let E : y2 = x3 +ax+b be an elliptic curve over a finite field 𝔽q of characteristic p ≥ 5. Let be an odd number, and S a subgroup of E(Fq) of order . Set S=S{OE}. Then a normalized separable isogeny ϕ:EE˜=E/S of degree can be written as

(1) ϕ(x,y)=N(x)D(x), yN(x)D(x),

where D(x) is the polynomial defined as


and N(x) is related to D(x) through the formula

(2) N(x)D(x)=xs(3x2+a)D(x)D(x)2(x3+ax+b)D(x)D(x).

Here T(x) denotes the derivative of a function T(x) and xQ the x-coordinate of a point QS*. With coefficients s, s2, s3 of the polynomial D(x), set v=a(1)+3(s22s2) and w = 3as +2b(−1)+5(s3 −3ss2 +3s3). Then the Weierstrass equation for E˜ is given by y2=x3+a˜x+b˜ with a˜=a5v$and$b˜=b7w. (Fast algorithms are shown also in [6] to compute the isogeny ϕ.)

Kernel polynomials

Partition the set S* into two parts S+ and S such that S=S+SandS={P:PS+}. We consider the kernel polynomial

(3) F(x)=QS+(xxQ)=xd+txd1+t2xd2++td

with d=12, and it satisfies D(x) = F(x)2. Schoof [30] studied the relation among coefficients t and ti’s of F(x). Instead of working with E˜, he works with the isomorphic curve Eˆ:y2=x3+aˆx+bˆ with aˆ=4a˜ and bˆ=6b˜. To E, we associate the reduced Weierstrass -function by (z)=1z2+k=1ckz2k with c1=a5 , c2=b7 and ck=3(k2)(2k+3)j=1k2cjck1j for k ≥ 3. Similarly, the function ˆ for Eˆ is defined using â and . Then the polynomial F(x) satisfies

(4) z1F((z))=exp12tz2k=1cˆkck(2k+1)(2k+2)z2k+2.

(This is by reduction of relations over C, and the characteristic p must be large enough to hold over 𝔽q.) From this equation, we can represent every coefficient tiwitht,ck's andcˆk's, and with t,a,b,a˜andb˜ since ck’s and cˆk's are obtained from a,b,a˜andb˜. (See [4, Chap. VII] for the first few coefficients of F(x).) In particular, every ti can be represented as an element of the multivariate polynomial ring Fq[t,a,b,a˜,b˜] when we regard all t,a,b,a˜,b˜ as variables.

2.3 Modular polynomials

For every integer ≥ 2, there exists the modular polynomial Φ(X, Y) ∈ ℤ[X, Y] to parameterize pairs of elliptic curves with a cyclic isogeny of degree between them. (See [32, Exercise 2.18, Chap. II].) For two elliptic curves EandE˜ over a finite field 𝔽q, there is an isogeny of degree from EtoE˜ with cyclic kernel if and only if Φ(j(E),j(E˜))=0. Every modular polynomial Φ(X, Y) is symmetric in each variable, and its integer coefficients become very large as increases. In particular, for a prime , the modular polynomial Φ(X, Y) is equal to the form

(5) X+1XY+Y+1+i,j,i+j<2aijXiYj

with aij ∈ ℤ, since there are precisely + 1 subgroups of the -torsion group of an elliptic curve E. (Such each subgroup corresponds to an isogenous curve of degree with a j-invariant, which is a zero of the polynomial Φ(X, j(E)).)

3 Algebraic approaches for isogeny problems

In this section, we propose several algebraic approaches. Before presenting our approaches, let us define our setting of isogeny problem as below:

Problem 1

Let =0e be a power of a small odd prime 0. Suppose that there exists an isogeny of degree between elliptic curves EandE˜ over a finite field 𝔽q. Given ,j=j(E)andȷ˜=j(E˜), find an isogeny ϕ:EE˜ of degree .

Compared to the general isogeny problem in [20], the isogeny degree is restricted to a prime power and it is given in our setting.Note that a solution of ϕ is not unique in general for the above problem. (In particular, the kernel of ϕ is not necessarily cyclic.) There are also many (compact) representations of a solution ϕ such as a chain of isogenies of low degrees, a sequence of j-invariants of intermediate curves, and a path in an isogeny graph between EandE~. The meet-in-the-middle approach is a practical way to solve isogeny problems. For the above isogeny problem, it builds two trees of isogenies of prime degree 0 from both the sides of E and E˜, respectively, and it finds a collision between the two trees to find the shortest path from E to E˜. While the meet-in-the-middle approach is a generic way for solving graph problems, we shall propose two algebraic approaches for solving the above isogeny problem.

3.1 First approach using modular polynomials

In this subsection, we present our first approach. Consider a chain of isogenies ϕk of prime degree 0 from E to E˜ as


Let jk denote the j-invariant of every elliptic curve Ek for 1 ≤ ke−1. In this approach, we regard j-invariants jk’s as variables, and consider a system of equations using modular polynomials (cf., recall that j and ȷ˜ are elements of 𝔽q.)

(6) Φ0(j,j1)=0,Φ0(jk,jk+1)=0(1ke2),Φ0(je1,ȷ˜)=0.

A solution of this system gives all j-invariants jk’s of intermediate curves Ek.

Here we propose a method to solve the system (6) efficiently by Gröbner basis algorithms. (See textbooks [2, 13] for Gröbner basis computation.) We assume that the exponent e of the isogeny degree is even with e=2e0 for a positive integer e0 for simple discussion. We divide the system (6) into two parts like the meet-in-the-middle approach. In terms of Gröbner basis computation, we consider two ideals in different multivariate polynomial rings


Both ideals I and Ĩ are zero-dimensional since j,ȷ˜Fq. In particular, the dimensions of 𝔽q-vector spaces Fq[j1,,je0]/I and Fq[je0,,je1]/I˜ are both at most (0+1)e0 due to the form of the modular polynomial (see Equation (5)). Moreover, the above generators give a Gröbner basis for the ideal I (resp., the ideal Ĩ) with the lex term order with respect to je0j2j1 (resp., je0je2je1). Then we can efficiently compute minimal polynomials g and g˜ of the variable je0 with respect to ideals I and Ĩ, respectively, by using the FGLM algorithm [17]. (In this case, simple linear algebra might be more efficient since degrees of g and g˜ are known. See Remark 3.1 and Appendix A.1 below for details.) By the GCD computation over the univariate polynomial ring Fq[je0], we obtain a common root of two minimal polynomials gandg˜. Such a common root gives a solution of je0.

Once a solution of je0 is found, the isogeny problem is divided into two isogeny problems of smaller degree 0e0= i.e., a divide-and-conquer strategy). By repeating this procedure, we can solve the whole isogeny problem.

Remark 3.1

Set g1=Φ0(j,j1), and let gk denote the minimal polynomial of the variable j k with respect to the ideal Φ0(j,j1),,Φ0(jk1,jk) for every 2 ≤ ke0. Putting Gk(X):=Φ0k(j,X), we have generically


for 3 ≤ ke0. (We verified by experiments that it holds in most cases.) For our target minimal polynomial g=ge0, it seems the best in performance to compute gk from gk−1 and Φ0(jk1,jk) recursively for 2 ≤ ke0, see Appendix A.1. In a similar way, we can compute another minimal polynomial ˜g efficiently.

The time complexity of the first approach shall be analyzed in Appendix A.1. The complexity depends on an algorithm for computing minimal polynomials, such as the FGLM algorithm and the GCD computation. Note that our first approach is not better than the meet-in-the-middle approach in time complexity.

Possible improvement (3-section method)

We introduce a possible improvement for the first approach. Our idea is to divide the system (6) into three parts. (cf., The original strategy is regarded as the “2-section method”.) With two parameters e1 and e2 satisfying 1 < e1 < e0 < e2 < e and e1ee2, consider two ideals in different multivariate polynomial rings


As in the 2-section method, we use the lex term order with je1j2j1 (resp., je2je2+1 je−1) for the zero-dimensional ideal I[1:e1] (resp., I˜[e2:e1]). Then the ideal I[1:e1] (resp., I˜[e2:e1]) includes a polynomial g(je1) (resp., g˜(je2)) such that the set of its roots contains the roots of Φ(j,je1) of level =0e1 (resp., Φ˜(je2,ȷ˜) of level ˜=0ee2). Namely, the polynomials gandg˜ are minimal polynomials for zero-dimensional ideals I[1:e1] and I˜[e2:e1], respectively. We then consider a new ideal


For this zero-dimensional ideal, we use the grevlex term order with je1je2je1+1je21je0 to find intermediate j-invariants from je1 to je2. With these j-invariants, we obtain the other j-invariants j1, . . . , je11 and je2+1,,je1 as in the 2-section method. The time complexity of this 3-section method might be improved so that it is better than that of the 2-section method.

3.2 Second approach using kernel polynomials

In this subsection, we present our second approach for solving the isogeny problem of prime power degree =0e. As in the first approach, we assume that the exponent e of the isogeny degree is even with e = 2e0 for simple discussion. In this approach, we take an intermediate elliptic curve E0 between EandE˜ such that there exist two normalized isogenies φandφ˜ of same degree =0e0 as


Let a,b(resp.,a˜,b˜) denote Weierstrass coefficients in 𝔽q defining the initial curve E (resp., the final curve E˜). These coefficients can be chosen with two j-invariants jandȷ˜ofEandE˜. In contrast, we use two variables a0, b0 as Weierstrass coefficients defining the intermediate curve E0. (That is, .) E0:y2=x3+a0x+b0.)

Like Equation (3), consider the kernel polynomial


obtained from the isogeny φ:EE0 with m=12. Now we regard t as an additional variable. Then we can represent the other coefficients t2, . . . , tm of F(x) as elements of the multivariate polynomial ring R = 𝔽q[t, a0, b0] by Schoof’s work [30], described in Subsection 2.2 (recall a, b ∈ 𝔽q). In other words, we regard the polynomial F(x) as an element of R[x]. Furthermore, like Equation (1), we can reconstruct the isogeny φ:EE0 with rational functions over R[x] as


for any point (x, y) ∈ E, by letting D(x) = F(x)2. Here the rational function T(x)=N(x)D(x)R(x) is given by Equation (2). Since φ(x, y) ∈ E0 for any point (x, y) ∈ E, we clearly have the relation

(7) (x3+ax+b)T(x)2=T(x)3+a0T(x)+b0.

By expanding this relation with respect to x, we obtain a set of equations S defined over R. (Every coefficient of xi in (7) corresponds to an equation in S.)

Similarly, we consider the kernel polynomial


obtained from another isogeny φ˜:E˜E0. By regarding t˜ as another variable, we can also regard F˜(x˜) as an element of the polynomial ring R˜[x˜]withR˜=Fq[t˜,a0,b0]$. We can also reconstruct the isogeny φ˜ as φ˜(x˜,y˜)=(T˜(x˜),y˜T˜(x˜)) for any point (x˜,y˜)E˜, for some rational function T˜(x˜)R˜(x˜). Similarly to the previous paragraph, the relation (x˜3+a˜x˜+b˜)T˜(x˜)2=T˜(x˜)3+a0T˜(x˜)+b0 gives another set of equations S˜ defined over the ring R˜.

Finally, we solve the system of equations in the union set SS˜ by using Gröbner basis algorithms over the multivariate polynomial ring Fq[t,t˜,a0,b0]. (We used the grevlex term order with tt˜a0b0 in our experiments.) A solution of (t,t˜,a0,b0) determines the Weierstrass equation for the intermediate curve E0, and also two kernels of isogenies φandφ~.

Remark 3.2

While the first approach requires many variables as the exponent e increases, this approach always requires four variables t,t˜,a0,b0. On the other hand, as e increases, total degrees of equations become large in this approach, while total degrees do not change in the first approach due to the use of the modular polynomial of same level 0. In contrast, one can consider several variants of this approach. As the simplest variant, we directly consider the kernel polynomial of the isogeny ϕ:EE˜ (without taking any intermediate curve between EandE~). This variant requires only one variable, but total degrees of equations become very large as e increases. On the other hand, we consider multiple intermediate curves as a variant. This variant requires many variables as the number of intermediate curves increases. As a result, the meet-in-the-middle type described in this subsection seems the best in performance from our preliminary experiments.

As in the first approach, the time complexity of the second approach shall be analyzed in Appendix A.2. In particular, the analysis shows that the second approach can never be asymptotically faster than the first one.

4 Experiments

In this section, we report experimental results of our algebraic approaches for solving the isogeny problem of prime power degrees (see Problem 1 for the problem).

4.1 Experiment parameters

For Problem 1, we fix parameters p = 2250 · 3159 − 1, q = p2 and E : y2 = x3 + x, extracted from SIKE-p503 parameters [21]. (The extension field 𝔽p2 is represented as 𝔽p[z]/(z2+1).) The initial curve E is a supersingular elliptic curve defined over 𝔽p2 having #E(Fp2)=(p+1)2=(22503159)2 and j = j(E) = 1728. We also take 3-powers = 3e (i.e., 0 = 3) as isogeny degrees for even exponents e=2e0. (cf., SIKE uses a combination of isogenies of degrees 2 and 3.) We follow [21] to generate the final supersingular curve E˜, isogenous to E of degree .

4.2 Implementation details

Here we describe details of implementation for our algebraic approaches and the meet-in-the-middle approach with Magma [5], a computational algebra system.

For our first approach by 2-section and 3-section methods, we used a combination of the modular polynomials ΦN(X, Y) for N = 3, 32, 33, which are pre-computed in Magma, in order to obtain the minimal polynomials g,g˜. For example, for = 310, we computed Gröbner bases for I=Φ33(j,j3),Φ32(j3,j5), I˜=Φ32(j5,j7),Φ33(j7,ȷ˜) with the Magma command GroebnerBasis to obtain g, g˜Fp2[j5], then computed the GCD of g, ˜g with the Magma commands GCD for the 2-section method.

For our second approach, we used Equation (4) to represent two kernel polynomials F(x)andF˜(x˜) of isogenies φ and φ˜, described in Subsection 3.2, as polynomials over Fq[t,t˜,a0,b0]. (As an alternative of (4), fast algorithms are introduced in [6] for computing the kernel polynomial of an isogeny, but they require very fast exponentiation.) We also used the grevlex term order with tt˜a0b0 to find a solution (t,t˜,a0,b0) from the union set of equations SS˜.

For the meet-in-the-middle approach, we construct two sets JandJ˜ of sequences (j, j1, . . . , je0 ) and (ȷ˜,ȷ˜1,,ȷ˜e0) of j-invariants of elliptic curves EkandE˜k, respectively. (Recall j = j(E) and ȷ˜=j(E˜).) Here Ek−1 and Ek (resp., E˜k1andE˜k) are isogenous of degree 3 for every 1 ≤ ke0, where we set E0=E(resp.,E˜0=E˜) for convenience. To construct sequences in JandJ˜, we use the modular polynomial of level 3. For example, we add each solution of Φ3(jk1,x) to a sequence (j, j1, . . . , jk−1) of length k. (We used the Magma command Roots to find a solution.) In constructing such sequences, we remove a sequence whose ending point already appeared in the other sequences, in order to reduce the sizes of two sets JandJ˜. In our experiments, it terminates when we find a pair of two sequences of JandJ˜ satisfying je0=ȷ˜e0 (i.e., a collision).

4.3 Experimental results

In Table 1, we summarize average running times of our two approaches and the meet-in-the-middle approach for solving the isogeny problem of degrees = 3e with even e from e = 6 up to 14. Specifically, we measured the running time of every approach until it finds the j-invariant or the Weierstrass coefficients of an intermediate curve between two isogenous curves EandE˜. We also experimented 5 times for every parameter set. All the experiments were performed using Magma 2.24-5 on 4.20 GHz Intel Core i7 CPU with 16 GByte memory. From Table 1, the first approach is the fastest for isogeny degrees up to = 310. In contrast, our second approach is costly and it did not terminate in one day for degrees larger than = 38. For degrees larger than = 310, the meet-in-the-middle approach is faster than our algebraic approaches. With respect to the memory usage, our first approach by the 2-section method requires about 64, 221 and 526MByte for = 310, 312 and 314, respectively, while the meet-in-the-middle approach requires about 24, 26 and 32 MByte for the same isogeny degrees.

Table 1

Average running times (seconds) of our two approaches and the meet-in-the-middle approach for the isogeny problem of degrees = 3e on supersingular elliptic curves over 𝔽p2 with 503-bit prime p = 2250 · 3159 − 1 (These parameters are from SIKE-p503 [21], and the initial curve E is given by y2 = x3 + x over 𝔽p2)

Isogeny degrees Our first approach Our second approach Meet-in-the-middle approach
= 3e 2-section 3-section
36 = 729 0.11 0.15 336.42 2.93
38 = 6561 0.19 0.45 > 1 day 9.61
310 = 59049 7.98 5.63 31.09
312 = 531441 287.77 159.39 96.75
314 = 4782969 5071.16 2725.01 292.82

5 Concluding remarks

In this paper, we proposed two algebraic approaches for solving isogeny problems of prime power degrees. The first one is a straightforward way using modular polynomials of small prime levels. The second one is a more complex way, which uses kernel polynomials of isogenies based on [30] and reconstructs the isogenies with Vélu’s formulae [6, 39]. From the analysis in Appendix A, our approaches are not asymptotically faster than the meet-in-the-middle approach. However, our experiments showed that the first approach is faster than the meet-in-the-middle approach for isogeny degrees up to = 310 over supersingular elliptic curves of SIKE-p503 [21] with a single classical computer. On the one hand, our first approach is applicable in the collision search step of the meet-in-the-middle approach. Such combination could make it faster in practice and reduce the memory size of the meet-in-the-middle approach. As future work, we would like to use the combination for solving isogeny problems of large degrees.


This work was supported by JSPS KAKENHI Grant Numbers 18H05836, 19K21026, and 19K22847, Japan


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A Time complexity analysis

In this appendix, we analyze time complexities of our algebraic approaches. Here we use the same notation as in Section 3. In this complexity analysis, we assume that and 0 are constants, and we set e and e0 as asymptotic parameters to estimate time complexities.

A.1 For the first approach

The first approach proceeds with the below two steps:

Step 1. Compute minimal polynomials gandg˜ of the variable je0 with respect to ideals IandI˜, respectively.

Step 2. By computing the GCD of gandg˜ over the univariate polynomial ring Fq[je0], we obtain a common root in 𝔽q of the minimal polynomials.

There are several methods to compute Step 1, such as the FGLM algorithm and the GCD computation. In the below, we estimate costs of the first approach using several different methods for Step 1.

Using the FGLM algorithm for Step 1

Applying the FGLM algorithm directly to the whole ideal

The cost of Step 1 by using the FGLM algorithm for the ideal I is estimated as O e0((0+1)e0)3=Oe0(0+1)3e0, since dimFqFq[j1,,je0]/I(0+1)e0 and the number of variables is e0. Similarly, the FGLM algorithm for another ideal Ĩ requires the same cost. Hence the complexity of Step 1 is

Using the normal form computation and simple linear algebra only

Instead of directly using of the FGLM algorithm, we can compute minimal polynomials gandg˜ based on the below well-known fact:

Lemma A.1

Let K be a field, I a zero-dimensional ideal of S = K [X1, . . . , Xn], and G a Gröbner basis of I with respect to a fixed term order. Let fS and g(t) ∈ K [t] with leading coefficient 1. We denote by mf (t) the minimal polynomial of f with respect to the ideal I. Then g(t) = mf (t) if and only if g(t) satisfies the following two conditions:

  1. NFI(g(f))=0, where NFI(g(f )) denotes the normal form of g(f) ∈ S with respect to G.

  2. For any h(t) ∈ K[t] with NFI(h(f))=0, we have deg(g) ≤ deg(h).

Recall from Remark 3.1 that we have generically g=Gk2Gkfor3ke0, where we set Gk(jk):=Φk(j,jk) for 1 ≤ ke0. (Until the end of this appendix, we assume this generic condition.) Since the degree of g equals to d:=deg(g)=deg(Ge02)+deg(Ge0)=(0+1)0e03+(0+1)0e01=O(0e0), it suffices to compute NFI(Xk) for 0kd=O(0e0) with X:=je0, and solve the linear equation NFI(Xd)+k=0d1ckNFI(Xk)=0 over 𝔽q, and then we have g=k=0dckXk by Lemma A.1. The cost of computing each normal form NFI(Xk) is O((0+1)2e0)=O(02e0). Hence the cost of computing NFI(Xk) for 0 ≤ kd is O((0e002e0))=O(03e0). The cost of solving the linear equation is not higher than O(03e0) since it is O(d3)=O(03e0) if one uses the Gaussian elimination naively. Similarly, computing g˜ requires O(03e0). Here we estimate that the cost of Step 1 is O(203e0)=O(03e0)=O(032e). In contrast, the cost of Step 2 mainly depends on the computation of the GCD of g and ̃g over 𝔽q. Note that the complexity of computing the GCD of two univariate polynomials of degree at most d over 𝔽q is O(d2) arithmetic over 𝔽q (if one uses classical polynomial arithmetic). Since both degrees of gandg˜ are equal to (0+1)0e03+(0+1)0e01=O(0e0), we estimate that the cost of Step 2 is O(02e0)=O(0e). Summing up, we roughly estimate that the complexity of the first approach is


if the FGLM algorithm is applied. Moreover, the complexity becomes


if only the normal form computation and simple linear algebra is used. For the parameter in our experiments (i.e., 0 = 3), the cost is O(e033e0) or O(33e0).

Using sequential GCD for Step 1

As mentioned in Remark 3.1, it is practical to compute minimum polynomials gk together with Gk(X) := Φ0k(j,X) sequentially for 2 ≤ ke0. (Note the target minimal polynomial g is given by ge0.) The minimal polynomial gk is the lowest degree polynomial in j k belonging to the ideal generated by Gk1(jk1)=Φ0k1(j,jk1) and Φ0(jk1,jk), that is


Once gk is computed, we obtain Gk as the quotient of gk divided by Gk−2 since gk(X) = Gk−2(X)Gk(X). Here we estimate the cost of computing Gk. Regarding Gk−1(jk−1) and Φ0(jk1,jk) as polynomials in jk−1 over 𝔽q[jk], the cost of computing G k can be upper-bounded by that of computing their greatest common divisor with respect to j k−1 by the Euclidean algorithm with pseudo division. This is due to the following reason; In the below steps (1) and (2), the target minimal polynomial can be computed by eliminating jk−1 from Gk1(jk1,jk) and Φ0(jk1,jk). The Euclidean algorithm with pseudo division corresponds to eliminating jk−1 by using leading coefficients as polynomials over 𝔽q[jk], whereas the Gröbner basis computation corresponds to eliminating jk−1 by using head terms as polynomials in 𝔽q[jk−1, j k] with respect to the lexicographical term order with jk−1jk. In each elimination step of the Euclidean algorithm with pseudo division, the elimination of a term is just the same as computing an S-polynomial (or its multiple by a monomial). To make our estimation simple, we can use so called the subresultant GCD algorithm which is the Euclidean algorithm with pseudo division and principal subresultant coefficient (PSC) and computes the resultant, that is, not the minimal polynomial but the characteristic polynomial (Gk2(jk))0Gk(jk) of j k with respect to Gk1(jk1),Φ0(jk1,jk).

Now we estimate the cost of computing the resultant of Gk−1 and Φ0(jk1,jk) with respect to jk−1 by the subresultant GCD, where we may assume that the remainder sequence is normal. Note that degjk1(Gk1(jk1))= (0+1)0k2=O(0k1), and degjk1(Φ0(jk1,jk))=0+1. The final target Gk is computed by conducting two procedures below:

  1. Using the division algorithm, compute Gk1(jk1,jk):=Gk1(jk1) mod Φ0(jk1,jk), where Φ0(jk1,jk) is a monic polynomial over 𝔽q[jk] of degree 0 + 1 in jk−1. Note that Gk1 is a univariate polynomial over 𝔽q[jk] in variable jk−1 of degree ≤ 0. More specifically, proceed with the following (we use simple notation A, B, C for polynomials and N for degrees):

    (1-1) Set AΦ0(jk1,jk), and BGk−1.

    (1-2) Set Ndegjk1(B), (B), and C ←(the coefficient of jk1N in B).

    (1-3) Compute Bjk1N(0+1) CA, and set BBjk1N(0+1) CA.

    (1-4) If N0, stop the loop. Otherwise set Ndegjk1 (B) and go back to (1-2).

    The resulting polynomial B coincides with the target Gk1.

  2. Compute the resultant of Gk1(jk1,jk) and Φ0(jk1,jk) by the subresultant GCD, and then recover Gk from the resultant.

We estimate the cost of the procedure (1). Let n be the number of times the loop executes, and let (Bs , Ns , Cs) for 1 ≤ sn be tuples of the successive values Bs of B, Ns of N and Cs of C respectively. We denote by Ms the maximum degree of the coefficients of jk1i for 0 ≤ iNs − 1 in Bs. Note that B1=Gk1Fq[jk1], N1=degjk1(Gk1)=O(0k1),C1Fq{0} and M1 ≤ 0. Recall from (2.5) that


and hence all the coefficients of jk1i for 0 ≤ iNs of jk1Ns(0+1)CsA are polynomials over 𝔽q[jk] of degrees at most deg(Cs)+0+1. From this, the degree of the coefficient of each jki with 0 ≤ iNs in Bs+1=Bsjk1Ns(0+1)CsA is at most max {Ms, deg(Cs)+0 +1}, and thus both deg(Cs+1) and Ms+1 are ≤ max {Ms, deg(Cs)+0 +1}. Since M1 ≤ 0 ≤ deg(C1)+0 +1, successively we have Ms ≤ (s −1)(0 +1) and deg(Cs)(s1)(0+1). Since A and Cs have at most (0+1)2+2=O((0+1)2) and deg(Cs) terms respectively, computing C sA requires O((s1)(0+1)3) arithmetic in 𝔽q, which is clearly dominant in the n-th loop. It follows from ndeg(B1)degjk1(A)+1= O(0k1) and s=1n(s1)(0+1)3=O(03n2) that the total cost of the procedure (1) is


By the estimation of the cost of computing a resultant of two polynomials, we estimate the complexity of the procedure (2). The procedure (2) requires O(02) arithmetic over 𝔽q[jk] since both Gk1 and Φ0(jk1,jk) have degree ≤ 0 + 1 = O(0). Since the degree of the resultant gˆk=(Gk2)0Gkisdˆk:=(0+1)20k2= O(0k), we may assume that all arithmetic over 𝔽q[jk] in the procedure (2) are arithmetic of two polynomials of degree O(00k)=O(0k) by considering pseudo division and PSC. Hence the complexity of the procedure (2) is estimated as O(02(0k)2)=O(02k) if one uses classical polynomial arithmetic, and O(02(0k)1+ϵ)=O(0(1+ϵ)k) for some 0 < ϵ < 1 if one uses fast polynomial computation (e.g., fast Fourier transform). Dividing the degree O(0k) -polynomial gˆk by (Gk2)0 provides Gk, and it is done in O02k (or O(0(1+ϵ)k)), which is equal to the cost of (2). Summing up, we estimate that the total cost is


or O(0(1+ϵ)e0)=O((1+ϵ)/2). For the parameter in our experiments (i.e., 0 = 3), the cost is O(32e0).

Remark A.1

For the meet-in-the-middle approach, we factorize the univariate polynomial Φ3(x, j) over 𝔽 p2 for the j-invariant j = j(E) of an elliptic curve E to find j-invariants of all elliptic curves 3-isogeneous to E. It costs roughly O(log p), and hence the total cost of the meet-in-the-middle approach is O(3e0logp). On the other hand, the total cost of our first approach is O(32e0) as discussed above, and hence we expect that our first approach could be faster than the meet-in-the-middle approach when log p>3e0. This is a main reason why our first approach is faster than the meet-in-the-middle approach in practice for small exponents e0.

A.2 For the second approach

We estimate the complexity of our second approach for solving the isogeny problem of prime degree =0e with e=2e0. We use the same notation as in Subsection 3.2. To estimate the complexity, we calculate the number of equations in the system, that is, #(SS˜), and the maximum value of the total degrees of the equations with respect to t,t˜,a0 and b0.

First, we bound the total degrees of tiandt˜i with respect to t,t˜,a0 and b0 for 2 ≤ im. The kernel polynomial F (x) for E is written as F=xm+txm1+k=0m2tmkxk. For each 2 ≤ im, the total degree of ti with respect to t, a0 and b0 is i, and thus it is bounded by m. Similarly, the total degree of t˜i with respect to t˜ a0 and b0 is bounded by m. Next, we explicitly write down a polynomial G=kckxkR[x] such that S={ck:0kdegxG}, where the coefficient ring is R = 𝔽q[t, a0, b0]. Recall that we obtain S by expanding (3.2) with respect to x. It follows from T=ND and D = F 2 that we have T=NF2NFF3, and thus (3.2) is rearranged as

(8) (x3+ax+b)NF2NFF32=N3F6+a0NF2+b0.

Note that


by (2.2), and hence N is of degree 2m + 1 with respect to x. We also note that each coefficient of N is a polynomial in R = 𝔽q[t, a0, b0] of total degree ≤ 2m + 1 = O(m). Multiplying by F 6 the both side of (8), we have


and set


which is of degree 6m+3 with respect to x. Writing G=k=06m+3ckxk with ckR for 0 ≤ k ≤ 6m+3, we clearly have S = {ck : 0 ≤ k ≤ 6m+3}, and the total degree of each cis equal to or less than 6m+3 = O(m). Similarly, we can construct a polynomial G˜=kc˜kxkR˜[x] with R˜=Fq[t˜,a0,b0] such that S˜={c˜k:0kdegxG˜} with degxG˜=6m+3. The total degree of each ̃ck is equal to or less than 6m + 3 = O(m).

Here we assume that we use the algorithm F 4, which is default for computing Gröbner bases in Magma, to solve a system of (non-homogeneous) multivariate equations. Let ISS˜ denote the ideal defined by SS˜. Let solv. deg(ISS˜) denote the solving degree of the ideal ISS˜Fq[t,t˜,a0,b0] with respect to the grevlex order. To estimate the complexity of computing a Gröbner basis of ISS˜, we use the following proposition proved by Caminata and Gorla:

Proposition A.1 ([7])

Let K be a field, and R = K [X1, . . . , Xn] the polynomial ring of n variables. Let f1, . . . , f be (not necessarily homogeneous) polynomials in R, and I the ideal generated by f1, . . . , f. We denote by s = solv. deg(I) the solving degree of I, that is, the highest degree of the polynomials involved in the computation of a Gröbner basis of I (we fix the grevlex order for Gröbner bases). Let di be the total degree of fi for 1 ≤ iℓ, and


Then the number of operations in K required to compute a Gröbner basis of I is


where 2 ≤ ω ≤ 3 is the exponent of matrix multiplication.

Since the total degree of each cSS˜is6m+3=O(m), we have

solv. deg(ISS˜)6m+3=O(m),

where the solving degree is not less than the maximum degree of input polynomials. It also follows from #(SS˜)2(6m+3)=12m+6=O(m) and

4+solv. deg(ISS˜)solv. deg(ISS˜)=4+solv. deg(ISS˜)4O(m4)

that the number of operations in 𝔽q required to compute a Gröbner basis of ISS˜ in the second approach is Om5=O()5=O05e0 by Proposition A.1, and so is the total complexity of the second approach. For the parameter in our experiments (i.e., 0 = 3), the complexity is not less than 35e0.

Received: 2019-06-05
Accepted: 2019-07-01
Published Online: 2020-11-17

© 2020 Y. Takahashi et al., published by De Gruyter

This work is licensed under the Creative Commons Attribution 4.0 International License.

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