
We give a new proof of Fürer's bound for the cost of multiplying bit integers in the bit complexity model. Unlike Fürer, our method does not require constructing special coefficient rings with “fast” roots of unity. Moreover, we establish the improved bound with . We show that an optimised variant of Fürer's algorithm achieves only , suggesting that the new algorithm is faster than Fürer's by a factor of . Assuming standard conjectures about the distribution of Mersenne primes, we give yet another algorithm that achieves .

Let denote the cost of multiplying two bit integers in the deterministic multitape Turing model [38] (commonly called “bit complexity”). Previously, the best known asymptotic bound for was due to Fürer [18, 19]. He proved that there is a constant such that
where , for , denotes the iterated logarithm, i.e.,
The main contribution of this paper is a new algorithm that yields the following improvement.
Fürer suggested several methods to minimise the value of in his algorithm, but did not give an explicit bound for . In section 7 of this paper, we outline an optimised variant of Fürer's algorithm that achieves . We do not know how to obtain using Fürer's approach. This suggests that the new algorithm is faster than Fürer's by a factor of .
The idea of the new algorithm is remarkably simple. Given two bit integers, we split them into chunks of exponentially smaller size, say around bits, and thus reduce to the problem of multiplying integer polynomials of degree with coefficients of bit size . We multiply the polynomials using discrete Fourier transforms (DFTs) over , with a working precision of bits. To compute the DFTs, we decompose them into “short transforms” of exponentially smaller length, say length around , using the Cooley–Tukey method. We then use Bluestein's chirp transform to convert each short transform into a polynomial multiplication problem over , and finally convert back to integer multiplication via Kronecker substitution. These much smaller integer multiplications are handled recursively.
The algorithm just sketched leads immediately to a bound of the form (1.1). A detailed proof is given in section 4. We emphasise that the new method works directly over , and does not need special coefficient rings with “fast” roots of unity, of the type constructed by Fürer. Optimising parameters and keeping careful track of constants leads to Theorem 1.1, which is proved in section 6. We also prove the following conditional result in section 9.
Conjecture 9.1 is a slight weakening of the Lenstra–Pomerance–Wagstaff conjecture on the distribution of Mersenne primes, i.e., primes of the form . The idea of the algorithm is to replace the coefficient ring by the finite field ; we are then able to exploit fast algorithms for multiplication modulo numbers of the form .
An important feature of the new algorithms is that the same techniques are applicable in other contexts, such as polynomial multiplication over finite fields. Previously, no Fürertype complexity bounds were known for the latter problem. The details are presented in the companion paper [24].
In the remainder of this section, we present a brief history of complexity bounds for integer multiplication, and we give an overview of the paper and of our contribution. More historical details can be found in books such as [21, Chapter 8].
Multiplication algorithms of complexity in the number of digits were already known in ancient civilisations. The Egyptians used an algorithm based on repeated doublings and additions. The Babylonians invented the positional numbering system, while performing their computations in base instead of . Precise descriptions of multiplication methods close to the ones that we learn at school appeared in Europe during the late Middle Ages. For historical references, we refer to [49, Section II.5] and [37, 5].
The first subquadratic algorithm for integer multiplication, with complexity , was discovered by Karatsuba [30, 31]. From a modern viewpoint, Karatsuba's algorithm utilises an evaluationinterpolation scheme. The input integers are cut into smaller chunks, which are taken to be the coefficients of two integer polynomials; the polynomials are evaluated at several wellchosen points; their values at those points are (recursively) multiplied; interpolating the results at those points yields the product polynomial; finally, the integer product is recovered by pasting together the coefficients of the product polynomial. This cuttingandpasting procedure is sometimes known as Kronecker segmentation (see section 2.6).
Shortly after the discovery of Karatsuba's algorithm, which uses three evaluation points, Toom generalised it so as to use evaluation points instead [51, 50], for any . This leads to the bound for fixed . Letting grow slowly with , he also showed that . The algorithm was adapted to the Turing model by Cook [10] and is now known as Toom–Cook multiplication. Schönhage obtained a slightly better bound [45] by working modulo several numbers of the form instead of using several polynomial evaluation points. Knuth proved that an even better complexity bound could be achieved by suitably adapting Toom's method [32].
The next step towards even faster integer multiplication was the rediscovery of the fast Fourier transform (FFT) by Cooley and Tukey [11] (essentially the same algorithm was already known to Gauss [27]). The FFT yields particularly efficient algorithms for evaluating and interpolating polynomials on certain special sets of evaluation points. For example, if is a ring in which is invertible, and if is a principal th root of unity (see section 2.2 for detailed definitions), then the FFT permits evaluation and interpolation at the points using only ring operations in . Consequently, if and are polynomials in whose product has degree less than , then the product can be computed using ring operations as well.
In [47], Schönhage and Strassen presented two FFTbased algorithms for integer multiplication. In both algorithms, they first use Kronecker segmentation to convert the problem to multiplication of integer polynomials. They then embed these polynomials into for a suitable ring and multiply the polynomials by using FFTs over . The first algorithm takes and , and works with finiteprecision approximations to elements of . Multiplications in itself are handled recursively, by treating them as integer multiplications (after appropriate scaling). The second algorithm, popularly known as the Schönhage–Strassen algorithm, takes where is a Fermat number. This algorithm is the faster of the two, achieving the bound . It benefits from the fact that is a principal th root of unity in , and that multiplications by powers of can be carried out efficiently, as they correspond to simple shifts and negations. At around the same time, Pollard pointed out that one can also work with where is a prime of the form , since then contains primitive th roots of unity [39] (although he did not give a bound for ).
Schönhage and Strassen's algorithm remained the champion for more than thirty years, but was recently superseded by Fürer's algorithm [18]. In short, Fürer managed to combine the advantages of the two algorithms from [47], to achieve the bound . Fürer's algorithm is based on the ingenious observation that the ring contains a small number of “fast” principal th roots of unity, namely the powers of , but also a large supply of much higherorder roots of unity inherited from . To evaluate an FFT over , he decomposes it into many “short” transforms of length at most , using the Cooley–Tukey method. He evaluates the short transforms with the fast roots of unity, pausing occasionally to perform “slow” multiplications by higherorder roots of unity (“twiddle factors”). A slightly subtle point of the construction is that we really need, for large , a principal th root of unity such that .
In [15] it was shown that the technique from [39] to compute modulo suitable prime numbers of the form can be adapted to Fürer's algorithm. Although the complexity of this algorithm is essentially the same as that of Fürer's algorithm, this method has the advantage that it does not require any error analysis for approximate numerical operations in .


Throughout the paper, integers are assumed to be handled in the standard binary representation. For our computational complexity results, we assume that we work on a Turing machine with a finite but sufficiently large number of tapes [38]. The Turing machine model is very conservative with respect to the cost of memory access, which is pertinent from a practical point of view for implementations of FFT algorithms. Nevertheless, other models for sequential computations could be considered [46, 20]. For practical purposes, parallel models might be more appropriate, but we will not consider these in this paper. Occasionally, for polynomial arithmetic over abstract rings, we will also consider algebraic complexity measures [8, Chapter 4].
In section 2, we start by recalling several classical techniques for completeness and later use: sorting and array transposition algorithms, discrete Fourier transforms (DFTs), the Cooley–Tukey algorithm, FFT multiplication and convolution, Bluestein's chirp transform, and Kronecker substitution and segmentation. In section 3, we also provide the necessary tools for the error analysis of complex Fourier transforms. Most of these tools are standard, although our presentation is somewhat ad hoc, being based on fixed point arithmetic.
In section 4, we describe a simplified version of the new integer multiplication algorithm, without any attempt to minimise the aforementioned constant . As mentioned in the sketch above, the key idea is to reduce a given DFT over to a collection of “short” transforms, and then to convert these short transforms back to integer multiplication by a combination of Bluestein's chirp transform and Kronecker substitution.
The complexity analysis of Fürer's algorithm and the algorithm from section 4 involves functional inequalities which contain postcompositions with logarithms and other slowly growing functions. In section 5, we present a few systematic tools for analysing these types of inequalities. For more information on this quite particular kind of asymptotic analysis, we refer the reader to [44, 16].
In section 6, we present an optimised version of the algorithm from section 4, proving in particular the bound (Theorem 1.1), which constitutes the main result of this paper. In section 7, we outline a similar complexity analysis for Fürer's algorithm. Even after several optimisations of the original algorithm, we were unable to attain a bound better than . This suggests that the new algorithm outperforms Fürer's algorithm by a factor of .
This speedup is surprising, given that the short transforms in Fürer's algorithm involve only shifts, additions and subtractions. The solution to the paradox is that Fürer has made the short transforms too fast. Indeed, they are so fast that they make a negligible contribution to the overall complexity, and his computation is dominated by the “slow” twiddle factor multiplications. In the new algorithm, we push more work into the short transforms, allowing them to get slightly slower; the quid pro quo is that we avoid the factor of two in zeropadding caused by Fürer's introduction of artificial “fast” roots of unity. The optimal strategy is actually to let the short transforms dominate the computation, by increasing the short transform length relative to the coefficient size. Fürer is unable to do this, because in his algorithm these two parameters are too closely linked. To underscore just how far the situation has been inverted relative to Fürer's algorithm, we point out that in our presentation we can get away with using Schönhage–Strassen for the twiddle factor multiplications, without any detrimental effect on the overall complexity.
We have chosen to base most of our algorithms on approximate complex arithmetic. Instead, following [39] and [15], we might have chosen to use modular arithmetic. In section 8, we will briefly indicate how our main algorithm can be adapted to this setting. This variant of our algorithm presents several analogies with its adaptation to polynomial multiplication over finite fields [24].
The question remains whether there exists an even faster algorithm than the algorithm of section 6. In an earlier paper [17], Fürer gave another algorithm of complexity under the assumption that there exist sufficiently many Fermat primes, i.e., primes of the form . It can be shown that a careful optimisation of this algorithm yields the bound . Unfortunately, odds are high that is the largest Fermat prime. In section 9, we present an algorithm that achieves the bound under the more plausible conjecture that there exist sufficiently many Mersenne primes (Theorem 1.2). The main technical ingredient is a variant of an algorithm of Crandall and Fagin [12] that permits efficient multiplication modulo , despite not being divisible by a large power of two.
It would be interesting to know whether the new algorithms could be
useful in practice. We have implemented an unoptimised version of the
algorithm from section 8 in the
Notations. We use Hardy's notations for , and for and . The symbol denotes the set of nonnegative real numbers, and denotes . We will write .
This section recalls basic facts on Fourier transforms and related techniques used in subsequent sections. For more details and historical references we refer the reader to standard books on the subject such as [2, 8, 21, 42].
In the Turing model, we have available a fixed number of linear tapes. An array of bit elements is stored as a linear array of bits. We generally assume that the elements are ordered lexicographically by , though this is just an implementation detail.
What is significant from a complexity point of view is that occasionally we must switch representations, to access an array (say 2dimensional) by “rows” or by “columns”. In the Turing model, we may transpose an matrix of bit elements in time , using the algorithm of [4, Appendix]. Briefly, the idea is to split the matrix into two halves along the “short” dimension, and transpose each half recursively.
We will also require more complex rearrangements of data, for which we resort to sorting. Suppose that is a totally ordered set, whose elements are represented by bit strings of length , and suppose that we can compare elements of in time . Then an array of elements of may be sorted in time using merge sort [33], which can be implemented efficiently on a Turing machine.
Let be a commutative ring with identity and let . An element is said to be a principal th root of unity if and
(2.1) 
for all . In this case, we define the discrete Fourier transform (or DFT) of an tuple with respect to to be where
That is, is the evaluation of the polynomial at .
If is a principal th root of unity, then so is its inverse , and we have
Indeed, writing , the relation (2.1) implies that
where if and otherwise.
Remark
Let be a principal th root of unity and let where . Then is a principal th root of unity and is a principal th root of unity. Moreover, for any and , we have
If and are algorithms for computing DFTs of length and , we may use (2.2) to construct an algorithm for computing DFTs of length as follows.
For each , the sum inside the brackets corresponds to the th coefficient of a DFT of the tuple with respect to . Evaluating these inner DFTs requires calls to . Next, we multiply by the twiddle factors , at a cost of operations in . (Actually, fewer than multiplications are required, as some of the twiddle factors are equal to . This optimisation, while important in practice, has no asymptotic effect on the algorithms discussed in this paper.) Finally, for each , the outer sum corresponds to the th coefficient of a DFT of an tuple in with respect to . These outer DFTs require calls to .
Denoting by the number of ring operations needed to compute a DFT of length , and assuming that we have available a precomputed table of twiddle factors, we obtain
For a factorisation , this yields recursively
The corresponding algorithm is denoted . The operation is neither commutative nor associative; the above expression will always be taken to mean .
Let be the butterfly algorithm that computes a DFT of length 2 by the formula . Then computes a DFT of length in time . Algorithms of this type are called fast Fourier transforms (or FFTs).
The above discussion requires several modifications in the Turing model. Assume that elements of are represented by bits.
First, for , we must add a rearrangement cost of to efficiently access the rows and columns for the recursive subtransforms (see section 2.1). For the general case , the total rearrangement cost is bounded by .
Second, we will sometimes use nonalgebraic algorithms to compute the subtransforms, so it may not make sense to express their cost in terms of . The relation (2.3) therefore becomes
where is the (Turing) cost of a transform of length over , and where is the cost of a single multiplication in .
Finally, we point out that requires access to a table of twiddle factors , ordered lexicographically by , for , . Assuming that we are given as input a precomputed table of the form , we must show how to extract the required twiddle factor table in the correct order. We first construct a list of triples , ordered by , in time ; then sort by in time (see section 2.1); then merge with the given root table to obtain a table , ordered by , in time ; and finally sort by in time . The total cost of the extraction is thus .
The corresponding cost for is determined as follows. Assuming that the table is given as input, we first extract the subtables of ()th roots of unity for in time . Extracting the twiddle factor table for the decomposition then costs ; the total over all is again .
Remark
Let be a principal th root of unity in and assume that is invertible in . Consider two polynomials and in . Let be the polynomial defined by
where the product of the DFTs is taken pointwise. By construction, we have , which means that for all . The product of and modulo also satisfies for all . Consequently, , , whence .
For polynomials with and , we thus obtain an algorithm for the computation of modulo using at most operations in . Modular products of this type are also called cyclic convolutions. If , then we may recover the product from its reduction modulo . This multiplication method is called FFT multiplication.
If one of the arguments (say ) is fixed and we want to compute many products (or cyclic convolutions) for different , then we may precompute , after which each new product can be computed using only operations in .
We have shown above how to multiply polynomials using DFTs. Inversely, it is possible to reduce the computation of DFTs — of arbitrary length, not necessarily a power of two — to polynomial multiplication [3], as follows.
Let be a principal th root of unity. For simplicity we assume that is even, and that there exists some with . Consider the sequences
Then , so for any we have
(2.5) 
Also, since is even,
Now let , and modulo . Then (2.5) implies that for all . In other words, the computation of a DFT of even length reduces to a cyclic convolution product of the same length, together with additional operations in . Notice that the polynomial is fixed and independent of in this product.
The only complication in the Turing model is the cost of extracting the in the correct order, i.e., in the order , given as input a precomputed table . We may do this in time by applying the strategy from section 2.3 to the pairs for . Similar remarks apply to the .
Remark
Multiplication in may be reduced to multiplication in using the classical technique of Kronecker substitution [21, Corollary 8.27]. More precisely, let and , and suppose that we are given two polynomials of degree less than , with coefficients and satisfying and . Then for the product we have . Consequently, the coefficients of may be read off the integer product where . Notice that the integers and have bit length at most , and the encoding and decoding processes have complexity .
The inverse procedure is Kronecker segmentation. Given and , and nonnegative integers and , we may reduce the computation of to the computation of a product of two polynomials of degree less than , and with and where . Indeed, we may cut the integers into chunks of bits each, so that , and . Notice that we may recover from using an overlapadd procedure in time . In our applications, we will always have , so that .
Kronecker substitution and segmentation can also be used to handle Gaussian integers (and Gaussian integer polynomials), and to compute cyclic convolutions. For example, given polynomials with , then for we have , so we may recover from the cyclic Gaussian integer product , where . In the other direction, suppose that we wish to compute for some . We may assume that the “real” and “imaginary” parts of and are nonnegative, and so reduce to the problem of multiplying , where and , and where the real and imaginary parts of are nonnegative and have at most bits.
In this section, we consider the computation of DFTs over in the Turing model. Elements of can only be represented approximately on a Turing machine. We describe algorithms that compute DFTs approximately, using a fixedpoint representation for , and we give complexity bounds and a detailed error analysis for these algorithms. We refer the reader to [7] for more details about multiple precision arithmetic.
For our complexity estimates we will freely use the standard observation that , since the multiplication of two integers of bit length reduces to multiplications of integers of bit length , for any fixed .
We will represent fixed point numbers by a signed mantissa and a fixed exponent. More precisely, given a precision parameter , we denote by the set of complex numbers of the form , where for integers and satisfying , i.e., . We write for the set of complex numbers of the form , where and ; in particular, for we always have . At every stage of our algorithms, the exponent will be determined implicitly by context, and in particular, the exponents do not have to be explicitly stored or manipulated.
In our error analysis of numerical algorithms, each is really the approximation of some genuine complex number . Each such comes with an implicit error bound ; this is a real number for which we can guarantee that . We also define the relative error bound for by . We finally denote by the “machine accuracy”.
Remark
In this section we give error bounds and complexity estimates for fixed point addition, subtraction and multiplication, under certain simplifying assumptions. In particular, in our DFTs, we only ever need to add and subtract numbers with the same exponent. We also give error bounds for fixed point convolution of vectors; the complexity of this important operation is considered later.
For , we define the “round towards zero” function by if and if . For , we define . Notice that and for any .
Proof. We have
and
whence .
Proof. We have
and
Consequently, .
Proposition 3.3 may be generalised to numerical cyclic convolution of vectors as follows.
Proof. Let denote the exact convolution, and write and . As in the proof of Proposition 3.3, we obtain and
The proof is concluded in the same way as Proposition 3.3.
Let and . Let be the branch of the square root function such that for . Using Newton's method [7, Section 3.5] and Schönhage–Strassen multiplication [47], we may construct a fixed point square root function , which may be evaluated in time , such that for all . For example, we may first compute some such that and , and then take ; the desired bound follows since .
Proof. The mean value theorem implies that where . For we have ; hence . By construction . We conclude that .
Proof. It suffices to compute . Starting from and , for each , we compute for using if and otherwise. Performing all computations with temporarily increased precision and corresponding , Lemma 3.5 yields . This also shows that the hypothesis is always satisfied, since . After rounding to bits, the relative error is at most .
A tight algorithm for computing DFTs of length is a numerical algorithm that takes as input an tuple and computes an approximation to the DFT of with respect to (or in the case of an inverse transform), such that
We assume for the moment that any such algorithm has at its disposal all necessary root tables with relative error not exceeding . Propositions 3.2 and 3.3 directly imply the following:
Proof. We have .
Proof. The inner and outer DFTs contribute factors of and , and by Proposition 3.3 the twiddle factor multiplications contribute a factor of . Thus
In this section we give the simplest version of the new integer multiplication algorithm. The key innovation is an alternative method for computing DFTs of small length. This new method uses a combination of Bluestein's chirp transform and Kronecker substitution (see sections 2.5 and 2.6) to convert the DFT to a cyclic integer product in for suitable .
Proof. Let , and suppose that we wish to compute the DFT of . Using Bluestein's chirp transform (notation as in section 2.5), this reduces to computing a cyclic convolution of suitable and . We assume that the and have been precomputed with .
We may regard and as cyclic polynomials with complex integer coefficients, i.e., as elements of . Write and , where with and . Now we compute the exact product using Kronecker substitution. More precisely, we have , so it suffices to compute the cyclic integer product , where . Then is the exact convolution of and , and rounding to precision yields in the sense of Proposition 3.4. A final multiplication by yields the Fourier coefficients .
To establish tightness, observe that and , so Proposition 3.4 yields where ; we conclude that . For , this means that the algorithm is tight; for , we may take .
For the complexity, observe that the product in reduces to three integer products of size . These have cost , and the algorithm also performs multiplications in , contributing the term.
Remark
Proof. We first reduce our integer product to a polynomial product using Kronecker segmentation (section 2.6). Splitting the two bit inputs into chunks of bits, we need to compute a product of polynomials with nonnegative bit coefficients and degrees less than . The coefficients of have bits, and we may deduce the desired integer product in time .
Let . To compute , we will use DFTs of length over , where . Zeropadding to obtain a sequence , and similarly for , we compute the transforms with respect to as follows.
Let and . Write with for and . We use the algorithm (see section 2.3), where for we take to be the tight algorithm for DFTs of length given by Proposition 4.1, and where is as in Corollary 3.9. In other words, we split the usual radix2 layers of the FFT into groups of layers, handling the transforms in each group with the Bluestein–Kronecker reduction, and then using ordinary Cooley–Tukey for the remaining layers.
We next compute the pointwise products , and then apply an inverse transform defined analogously to . A final division by (which is really just an implicit adjustment of exponents) yields approximations .
Since and are tight by Propositions 3.8, 4.1 and Corollary 3.9, we have , and similarly for . Thus , so after the inverse transform (since for ). In particular, , so we obtain the exact value of by rounding to the nearest integer.
Now we analyse the complexity. Using Proposition 3.6, we first compute a table of roots in time , and then extract the required twiddle factor tables in time (see section 2.3). For the Bluestein reductions, we may extract a table of th roots in time , and then rearrange them as required in time (see section 2.5). These precomputations are then all repeated for the inverse transforms.
By Corollary 3.9, Proposition 4.1 and (2.4), each invocation of (or ) has cost
The cost of the pointwise multiplications is subsumed within this bound.
It is now a straightforward matter to recover Fürer's bound.
Proof. Let for . By Theorem 4.3, there exists and such that
for all . Let for , . Increasing if necessary, we may assume that for , so that the function is welldefined. Increasing if necessary, we may also assume that for all .
We prove by induction on that for all . If , then , so the bound holds. Now suppose that . Since , we have , so by induction .
Finally, since , we have , so for .
This section is devoted to developing a framework for handling recurrence inequalities, similar to (4.1), that appear in subsequent sections.
Let be a smooth increasing function, for some . We say that is an iterator of if is increasing and if
for all sufficiently large .
For instance, the standard iterated logarithm defined in (1.2) is an iterator of . An analogous iterator may be defined for any smooth increasing function for which there exists some such that for all . Indeed, in that case,
is welldefined and satisfies (5.1) for all . It will sometimes be convenient to increase so that is satisfied on the whole domain of .
We say that is logarithmically slow if there exists an such that
for . For example, the functions , , and are logarithmically slow, with respectively.
Proof. The case is clear. For , let . By induction for large , so for large .
In this paper, the main role played by logarithmically slow functions is to measure size reduction in multiplication algorithms. In other words, multiplication of objects of size will be reduced to multiplication of objects of size , where for some logarithmically slow function . The following result asserts that, from the point of view of iterators, such functions are more or less interchangeable with .
Proof. First consider the case where in (5.2), i.e., assume that for some constant and all . Increasing and if necessary, we may assume that for all , and that .
We claim that
for all . Indeed, if , then
Now, given any , let , so . For any we have , so fold iteration of (5.3), starting with , yields
Moreover this shows that for , so . Since and , we obtain .
Now consider the general case . Let , so that is an iterator of . By the above argument , and so .
The next result, which generalises and refines the argument of Theorem 4.4, is our main tool for converting recurrence inequalities into actual asymptotic bounds for solutions. We state it in a slightly more general form than is necessary for the present paper, anticipating the more complicated situation that arises in [24].
Let and . Let , and let be any function satisfying the following recurrence. First, for all , . Second, for all , , there exist with , and weights with , such that
Then we have for all , .
Proof. Let , , and be as above. Define for . We claim that there exists , depending only on and , such that
for all . Indeed, let . First suppose , so that . For any , we have , so
and hence . This last inequality also clearly holds if (since ). By Lemma 5.2 we obtain .
Define a sequence of real numbers by the formula
We claim that
for all . Indeed, let . If then (5.5) holds as . If then by (5.4), so and hence .
Now let . We will prove by induction on that
for all . The base case , i.e., , holds by assumption. Now assume that , so . By hypothesis there exist , , and with , such that
Since , we obtain
Finally, the infinite product
certainly converges, so we have for . Setting , by (5.4) we obtain for all .
In this section, we present an optimised version of the new integer multiplication algorithm. The basic outline is the same as in section 4, but our goal is now to minimise the “expansion factor” at each recursion level. The necessary modifications may be summarised as follows.
Since Bluestein's chirp transform reduces a DFT to a complex cyclic convolution, we take the basic recursive problem to be complex cyclic integer convolution, i.e., multiplication in , rather than ordinary integer multiplication.
In multiplications involving one fixed operand, we reuse the transform of the fixed operand.
In a convolution of length with input coefficients of bit size , the size of the output coefficients is , so the ratio of output to input size is . We increase from to , so as to reduce the inflation ratio from to .
We increase the “short transform length” from to . The complexity then becomes dominated by the Bluestein–Kronecker multiplications, while the contribution from ordinary arithmetic in becomes asymptotically negligible. (As noted in section 1, this is precisely the opposite of what occurs in Fürer's algorithm.)
We begin with a technical preliminary. To perform multiplication in efficiently using FFT multiplication, we need to be divisible by a high power of two. We say that an integer is admissible if , where (note that for all ). We will need a function that rounds a given up to an admissible integer. For this purpose we define for . Note that may be computed in time .
Proof. We have , which implies (6.1). Since and , we have and thus , i.e., . In particular , so is admissible. (In fact, one easily checks that is the smallest admissible integer ).
Remark
Now let be admissible, and consider the problem of computing products with , i.e., products with one fixed operand. Denote the cost of this operation by . Our algorithm for this problem will perform forward DFTs and inverse DFTs, so it is convenient to introduce the normalisation
This is welldefined since clearly . Roughly speaking, may be thought of as the notional cost of a single DFT.
The problem of multiplying bit integers may be reduced to the above problem by using zeropadding, i.e., by taking and . Since and , we obtain . Thus it suffices to obtain a good bound for .
The recursive step in the main multiplication algorithm involves computing “short” DFTs via the Bluestein–Kronecker device. As pointed out in section 2.5, this leads to a cyclic convolution with one fixed operand. To take advantage of the fixed operand, let denote the cost of computing independent DFTs of length over , and let . Then we have the following refinement of Proposition 4.1. As usual we assume that the necessary Bluestein root table has been precomputed.
Proof. We use the same notation and algorithm as in the proof of Proposition 4.1, except that in the Kronecker substitution we take , so that the resulting integer multiplication takes place in . The proof of tightness is identical to that of Proposition 4.1 (this is where we use the assumption ). For the complexity bound, note that is admissible by construction, so for any we have Here we have used the fact that is fixed over all these multiplications. Dividing by and taking suprema over yields the result.
The next result gives the main recurrence satisfied by (compare with Theorem 4.3).
Proof. Let be admissible and sufficiently large, and consider the problem of computing products , for . Let , so that , and let .
We cut the inputs into chunks of size , i.e., if is one of the inputs, we write , where , and where the real and imaginary parts of have absolute value at most . Thus , and for any we may encode as a polynomial .
We will multiply the desired (cyclic) polynomials by using DFTs of length over where . We construct the DFTs in a similar way to section 4. Let and . Write with for and . We use the tight algorithm , where for we take to be the tight algorithm for DFTs of length given by Proposition 6.3, and where is as in Corollary 3.9. Thus, for the first groups of layers, we use Bluestein–Kronecker to reduce to complex integer convolution of size , and the remaining layers are handled using ordinary Cooley–Tukey. We write for the analogous inverse transform.
To check the hypothesis of Proposition 6.3, we observe that for sufficiently large , as is divisible by where , and
Denote by the cost of a single invocation of (or ). By Corollary 3.9 and (2.4), we have
The last term is the rearrangement cost, and simplifies to . The second term covers the invocations of , and simplifies to , so is absorbed by the term. The first term covers the invocations of . By definition , and since , Proposition 6.3 yields
Thus
We will use Schönhage–Strassen's algorithm for fixed point multiplications in . Since , we may take . Thus the term becomes
(We could of course use our algorithm recursively for these multiplications; however, it turns out that Schönhage–Strassen is fast enough, and leads to simpler recurrences. In fact, the algorithm asymptotically spends more time rearranging data than multiplying in !)
Since , and since , by Lemma 6.1 we have
We also have and , so
Thus
and consequently
To compute the desired products, we must execute forward transforms and inverse transforms. For each product, we must also perform pointwise multiplications in , at cost . As in the proof of Theorem 4.3, the cost of all necessary root table precomputations is also bounded by . Thus we obtain
Dividing by and taking suprema yields the bound (6.2).
The error analysis is almost identical to the proof of Theorem 4.3, the only difference being that is replaced by . Denoting one of the products by , we have exactly as in Theorem 4.3. Thus , and again we obtain by rounding to the nearest integer.
Finally we show how to define . We already observed that . Thus there exists a constant such that for large , so we may take .
Now we may prove the main theorem announced in the introduction.
Proof of Theorem 1.1. Let and be as in Theorem 6.4. Increasing if necessary, by Lemma 5.1 we may assume that for , and that .
Let for admissible . By the theorem, there exist constants such that for all admissible , there exists an admissible with
Increasing if necessary, we may also assume that for all admissible . Taking to be the set of admissible integers, we apply Proposition 5.3 with , , , and for each admissible setting , , and as above. We conclude that , and hence as runs over admissible integers. We already pointed out that .
As pointed out in the introduction, Fürer proved that for some , but did not give an explicit bound for . In this section we sketch an argument showing that one may achieve in Fürer's algorithm, by reusing tools from previous sections, especially section 6.
At the core of Fürer's algorithm is the ring , which contains the principal th root of unity . Note that is a direct sum of copies of , and hence not a field (for ). A crucial observation is that is a “fast” root of unity, in the sense that multiplication by and its powers can be achieved in linear time, as in Schönhage–Strassen's algorithm. For any , we need to construct a th root of , which is itself a th principal root of unity. We recall Fürer's construction of as follows.
is a principal th root of unity with . The coefficients of have absolute value .
Proof. See [19, Section 4].
As our basic recursive problem, we will consider multiplication in , where is divisible by a high power of two. We will refer to the last property as “admissibility”, but we will not define it precisely. We write for the cost of such products with one fixed argument, and for the normalised cost, exactly as in section 6.
Fürer worked with rather than , but, since we are interested in constant factors, and since the recursive multiplication step involves multiplication of complex quantities, it simplifies the exposition to work systematically with complexified objects everywhere.
For suitable parameters and , we will encode elements of as (nega)cyclic polynomials in , where as above. We choose the parameters later; for now we require only that divides and that (so that the coefficients are not too small).
The encoding proceeds as follows. Given , we split into parts of bits. Each is cut into even smaller pieces of bits. Then is encoded as
and an element is encoded as . (Notice that the coefficients of are zero for ; this zeropadding is the price Fürer pays for introducing artificial roots of unity.)
We represent complex coefficients by elements of for a suitable precision parameter . The exponent varies during the algorithm, as explained in [19]; nevertheless, additions and subtractions only occur for numbers with the same exponent, as in the algorithms from sections 4 and 6.
Given , to successfully recover the product from the polynomial product , we must choose , where is an allowance for numerical error. Certainly , and, as shown by Fürer, we may also take (an analogous conclusion is reached in sections 4 and 6). Thus we may assume that .
We must now show how to compute a product , for . Fürer handles these types of multiplications using “halfDFTs”, i.e., DFTs that evaluate at odd powers of , where is a principal th root of unity such that (Lemma 7.1). To keep terminology and notation consistent with previous sections, we prefer to make the substitution , i.e., writing , we put , and similarly for and . This reduces the problem to computing the product in . The change of variable imposes a cost of , where is the cost of a multiplication in .
So now consider a product , where . Let , so that . Let , and write with for and . For each , let be the algorithm for DFTs of length that applies the usual Cooley–Tukey method, taking advantage of the fast th root of unity The complexity of is , since it performs lineartime operations on objects of bit size . Let be the complexity of the algorithm for DFTs of length over . Then (2.4) yields
The first term is bounded by , since .
Let us now consider the second term , which describes the cost of the twiddle factor multiplications. This term turns out to be the dominant one. Both Kronecker substitution and FFT multiplication may be considered for multiplication in , but it turns out that Kronecker substitution is faster (a similar phenomenon was noted in Remark 4.2). So we reduce multiplication in to multiplication in where is admissible and divisible by . For any reasonable definition of admissibility we then have , provided that is somewhat smaller than . (In the interests of brevity, we will not specify the terms for the remainder of the argument. They can all be controlled along the lines of section 6.) Most of the twiddle factors are reused many times, so we will assume that , where the factor counts the two (rather than three) DFTs needed for each multiplication of size . The term of interest then becomes
Since and , this yields
To minimise the leading constant, we must choose to grow faster than , and to grow faster than . For example, taking and leads to and . The function mapping to is then bounded by a logarithmically slow function, and a similar argument to section 6 shows that .
Shortly after Fürer's algorithm appeared, De et al [15] presented a variant based on modular arithmetic that also achieves the complexity bound for some . Roughly speaking, they replace the coefficient ring with the field of adic numbers, for a suitable prime . In this context, working to “finite precision” means performing computations in , where is a precision parameter.
The main advantage of this approach is that the error analysis becomes trivial; indeed is a ring (unlike our ), and arithmetic operations never lead to precision loss (unless one divides by , which never happens in these algorithms). The main disadvantage is that there are certain technical difficulties associated with finding an appropriate ; this is discussed in section 8.2 below.
The aim of this section is to sketch an analogue of the algorithm of section 6 that achieves using modular arithmetic instead of . We assume familiarity with adic numbers, referring the reader to [22] for an elementary introduction.
For the basic problem, we take multiplication in , where is admissible (in the sense of section 6) and where one of the arguments is fixed over multiplications. As before, we take , and cut the inputs into chunks of bits. Thus we reduce to multiplying polynomials in with coefficients of at most bits. The coefficients of the product have at most bits.
Let be a prime such that , so that contains a primitive th root of unity . The problem of finding such and is discussed in the next section; for now we assume only that . We may then embed the multiplication problem into , and use DFTs with respect to to compute the product. On a Turing machine, we cannot represent elements of exactly, so we perform all computations in where
This choice ensures that , so knowledge of the product in determines it unambiguously in .
To compute each DFT, we first use the Cooley–Tukey algorithm to decompose it into “short transforms” of length , where . (As in section 6, there are also residual transforms of length for some , whose contribution to the complexity is negligible.) Multiplications in , such as the multiplications by twiddle factors, are handled using Schönhage–Strassen's algorithm, with the divisions by being reduced to multiplication via Newton's method. We then use Bluestein's algorithm to convert each short transform to a cyclic convolution of length over , and apply Kronecker substitution to convert this to multiplication in , where is the smallest admissible integer exceeding . This multiplication is then handled recursively.
Now, since , , and , we have , and hence , just as in section 6. The rest of the complexity analysis follows exactly as in the proof of Theorem 6.4, except for the computation of and , which is considered below.
Remark
Given a transform length for , our aim is to find a prime such that , i.e., such that divides . Denote by the smallest such prime.
HeathBrown has conjectured that [25], but given the current state of knowledge in number theory, we are only able to prove a result of the following type.
Proof. This is a special case of Linnik's theorem [34, 35], which states that there exist constants and such that for any with , there exists a prime number with . The best currently known estimate for is due to Xylouris [53]. Applying this result for and , we get the bound for large enough . The complexity bound follows by testing for primality until we find , using a polynomial time primality test [1].
The difficulty with this result — already noted in [15] — is that the time required to find greatly exceeds the time bound we are trying to prove for !
To avoid this problem, De et al suggested using a multivariate splitting, i.e., by encoding each integer as a polynomial in for suitable , say . One then uses dimensional DFTs to multiply the polynomials. Since the transform length is shorter, one can get away with a smaller . Unfortunately, this introduces further zeropadding and leads to a larger value of , ruining our attempt to achieve the bound .
On the other hand, we note that the problem only really occurs at the top recursion level. Indeed, at deeper recursion levels, there is exponentially more time available at the previous level to compute . So one possible workaround is to use a different, sufficiently fast algorithm at the top level, such as Fürer's algorithm, and then switch to the algorithm sketched in section 8.1 for the remaining levels. In this way one still obtains the bound , and asymptotically almost all of the computation is done using the algorithm of section 8.1.
If one insists on avoiding entirely, there are still many choices: one could use the algorithm of De et al at the top level, or use a multivariate version of the algorithm of section 8.1. One could even use the Schönhage–Strassen algorithm, whose main recursive step yields the bound ; applying this three times gives , and then to multiply integers with bits, one can find a suitable prime using Lemma 8.2 in time .
Another way to work around the problem is to assume the generalised Riemann hypothesis (GRH). De et al pointed out that under GRH, it is possible to find a suitable prime efficiently using a randomised algorithm. Here we show that, under GRH, we can even use deterministic algorithms.
Proof. The first bound is given in [26], and the complexity bound follows similarly to the proof of Lemma 8.2.
To use this result, we must modify the algorithm of section 8.1 slightly. Choose a constant so that we can compute in time , as in Lemma 8.3. Increase the coefficient size from to , and change the definition of admissibility accordingly. The transform length then decreases to , and the cost of computing decreases to only . The rest of the complexity analysis is essentially unchanged; the result is an algorithm with complexity , working entirely with modular arithmetic, in which the top recursion level does not need any special treatment.
Finally, we consider the computation of a suitable approximation to a th root of unity in .
Proof. We may find a generator of deterministically in time [48]. Then is a primitive th root of unity in , and there is a unique primitive th root of unity congruent to modulo . Given , we may compute using fast Newton lifting in time [9, Section 12.3].
In the context of section 8.1, we may assume that and , so the cost of finding is . This is certainly less than the cost of finding itself, using either Lemma 8.2 or Lemma 8.3.
It is natural to ask whether the approaches from sections 6, 7 or 8 can be further optimised, to obtain a complexity bound with .
In Fürer's algorithm, the complexity is dominated by the cost of multiplications in . If we could use a similar algorithm for a much simpler , then we might achieve a better bound. Such an algorithm was actually given by Fürer [17], under the assumption that there exist sufficiently many Fermat primes, i.e., primes of the form . More precisely, his algorithm requires that there exists a positive integer such that for every , the sequence contains a prime number. The DFTs are then computed directly over for suitable , taking advantage of the fact that contains a fast th primitive root of unity (namely the element 2) as well as a th primitive root of unity. It can be shown that a suitably optimised version of this hypothetical algorithm achieves : we still pay a factor of two due to the fact that we compute both forward and inverse transforms, and we pay another factor of two for the zeropadding in the recursive reduction. Unfortunately, it is likely that is the last Fermat prime [13].
In the algorithm of section 6, a potential bottleneck arises during the short transforms, when we use Kronecker substitution to multiply polynomials in . We really only need the high bits of each coefficient of the product (i.e., of the real and imaginary parts), but we are forced to allocate roughly bits per coefficient in the Kronecker substitution, and then we discard roughly half of the output. This problem is similar to the wellknown obstruction that prevents us from using FFT methods to compute a “short product”, i.e., the high bits or low bits of the product of two bit integers, any faster than computing the full bits.
In this section, we present a variant of the algorithm of section 6, in which the coefficient ring is replaced by a finite field , where is a Mersenne prime. Thus “short products” are replaced by “cyclic products”, namely by multiplications modulo . This saves a factor of two at each recursion level, and consequently reduces from to .
This change of coefficient ring introduces several technical complications. First, it is of course unknown if there are infinitely many Mersenne primes. Thus we are forced to rely on unproved conjectures about the distribution of Mersenne primes.
Second, is always prime (except possibly at the top recursion level). Thus we cannot cut up an element of into equalsized chunks with an integral number of bits, and still expect to take advantage of cyclic products. In other words, is very far from being admissible in the sense of section 6. To work around this, we deploy a variant of an algorithm of Crandall and Fagin [12], which allows us to work with chunks of varying size. The Crandall–Fagin algorithm was originally presented over , and depended crucially on the fact that contains suitable roots of . In our setting, we work over , where is a Mersenne prime exponentially smaller than . Happily, contains suitable roots of , and this enables us to adapt their algorithm to our setting. Moreover, since , the field contains roots of unity of high poweroftwo order, namely of order , so we can perform FFTs over very efficiently.
Finally, we can no longer use Kronecker substitution, as this would reintroduce the very zeropadding we are trying to avoid. Instead, we take our basic problem to be polynomial multiplication over (where is not necessarily prime). After the Crandall–Fagin splitting step, we have a bivariate multiplication problem over , which is solved using 2dimensional FFTs over . These FFTs are in turn reduced to 1dimensional FFTs using standard methods; this dimension reduction is, roughly speaking, the analogue of Kronecker substitution in this algorithm. (Indeed, it is also possible to give an algorithm along these lines that works over but avoids Kronecker substitution entirely; this still yields because of the “short product” problem mentioned above.) For the 1dimensional transforms, we use the same technique as in previous sections: we use Cooley–Tukey's algorithm to decompose them into “short transforms” of exponentially shorter length, then use Bluestein's method to convert them to (univariate) polynomial products, and finally evaluate these products recursively.
Let denote the number of Mersenne primes less than . Based on probabilistic arguments and numerical evidence, Lenstra, Pomerance and Wagstaff have conjectured that
as , where is the Euler constant [52, 40]. Our fast multiplication algorithm relies on the following slightly weaker conjecture.
Proof. The required prime exists since for we have
An integer of the form may be tested for primality in time using the Lucas–Lehmer primality test [13]. A simple way to compute is to apply this test successively for all ; this takes time . A primitive th root of unity may be computed by the formula in time ; see [43] or [14, Corollary 5].
Let be a Mersenne number (not necessarily prime). The main integer multiplication algorithm depends on a variant of Crandall and Fagin's algorithm that reduces multiplication in to multiplication in , where is a suitably smaller Mersenne prime (assuming that such a prime exists).
To explain the idea of this reduction, we first consider the simpler univariate case, in which we reduce multiplication in to multiplication in . Here we require that , that and that . For any , we will write and .
Assume that we wish to compute the product of . Considering and as elements of modulo , we decompose them as
(9.1) 
where
We regard and as complex “digits” of and , where the base varies with the position . Notice that takes only two possible values: or .
For , let
so that . For any , define as follows. Choose so that lies in the interval , and put
From (9.2), we have
Since the left hand side lies in the interval , this shows that . Now, since and , we have
where
Since and similarly for , we have . Note that we may recover from in time , by a standard overlapadd procedure (provided that ).
Let be the inverse of modulo ; this inverse exists since we assumed . Let , so that
since has order in . The quantity plays the same role as the real th root of appearing in Crandall–Fagin's algorithm.
Now define polynomials by and for , and let be their (cyclic) product. Then
coincides with the reinterpretation of as an element of . Moreover, we may recover unambiguously from , as and . Altogether, this shows how to reduce multiplication in to multiplication in .
Remark
Generalising the discussion of the previous section, we now show how to reduce multiplication in , for a given , to multiplication in . For this, we require that , that and that .
Indeed, consider two cyclic polynomials and in . We cut each of the coefficients into chunks and of bit size at most , using the same varying base strategy as above. With and as before, we next form the bivariate cyclic polynomials
in . Setting
the same arguments as in the previous section yield
Using the assumption that , we recover the coefficients , and hence the product , from the bivariate cyclic convolution product .
Let and (not necessarily prime). We will take our basic recursive problem to be multiplication in for suitable . We need somewhat larger than ; this is analogous to the situation in section 6, where we chose a “short transform length” somewhat larger than the coefficient size. Thus we set where is defined as follows.
(9.3) 
for all , and such that we may compute in time .
Proof. Let . Using [6], we may construct a function such that for all , and which may be computed in time . One checks that for all , so for . Thus is increasing, and has the desired properties.
We say that an integer is admissible if it is of the form where for some . (This should not be confused with the notion of admissibility of section 6.) An element of is then represented by bits. Note that is strictly increasing, so there is a onetoone correspondence between integers and admissible . For we define to be the smallest admissible integer .
Proof. From (9.3) we have ; this immediately implies that .
Suppose that we wish to compute for some . We assume that is large enough that the definition makes sense and so that . One checks that , so and hence . To find the smallest suitable , we may simply compute for each , and compare with . This takes time .
Now let , and . Consider the problem of computing products with . We denote by the complexity of this problem, where is the admissible integer corresponding to . As in section 6, we define .
Notice that multiplication of two integers of bit size reduces to the above problem, for , via a suitable Kronecker segmentation. Indeed, let for some , and encode the integers as integer polynomials of degree less than with coefficients of bit size . The desired product may be recovered from the product in , as
for large . Thus, as in section 6, we have , and it suffices to obtain a good bound for .
Now suppose additionally that is prime. In this case is a field, and as noted above, it contains th roots of unity, so we may define DFTs of length over for any . In particular, for we may use Bluestein's algorithm to compute DFTs of length . Denote by the cost of evaluating independent DFTs of length over , and put . Here we assume as usual that a th root of unity is known, and that the corresponding Bluestein root table has been precomputed.
Let us apply these definitions in the case ; this is permissible, as for sufficiently large . Since convolution of length over is exactly the basic recursive problem, and since one of the operands is fixed, we have , where , and hence
Proof. Let with . Assume that we wish to compute products with one fixed operand. Our goal is to reduce to a problem of the same form, but for exponentially smaller .
Choose parameters. Let be the smallest Mersenne prime larger than . By Proposition 9.2, we have , whence , for some absolute constant . Moreover, we may compute , together with a primitive th root of unity in , in time . We define and .
The algorithm must perform various multiplications in , at cost . For simplicity we will use Schönhage–Strassen's algorithm for these multiplications, i.e., we will take . Since , we have
Crandall–Fagin reduction. We use the framework of section 9.3 to reduce the basic multiplication problem in to multiplication in for suitable . We take where
We also write . The definition of makes sense for large since . Let us check that the hypotheses of section 9.3 are satisfied for large . We have and hence ; in particular, , so , and also . Since , we also have , and thus since .
We also note for later use the estimate
Indeed, since we have
and we already noticed earlier that .
To assess the cost of the Crandall–Fagin reduction, we note that computing the and costs (see Remark 9.3), the splitting itself and final overlapadd phase require time , and the various multiplications by , and have cost .
Reduction to poweroftwo lengths. Next we reduce multiplication in to multiplication in , where . In fact, since , these rings are isomorphic, via the map that sends to and to . Evaluating this isomorphism corresponds to rearranging the coefficients according to the rule , where is the exponent of and where and are the exponents of and . This may be achieved in time using the same sorting strategy as in section 2.3. The inverse rearrangement is handled similarly.
Reduction to univariate transforms. For multiplication in , we will use bivariate DFTs over . This is possible because contains both th and th primitive roots of unity, namely and , since and . More precisely, we must perform forward bivariate DFTs and inverse bivariate DFTs of length over , and multiplications in . Each bivariate DFT reduces further to univariate DFTs of length over (with respect to ) and univariate DFTs of length over (with respect to ). Interspersed between these steps are various matrix transpose operations of total cost , to enable efficient access to the “rows” and “columns” (see section 2.1).
Multiplications in are handled by zeropadding, i.e., we first use Cooley–Tukey to multiply in , and then reduce modulo . The total cost of these multiplications is .
Reduction to short transforms. Consider one of the “long” univariate DFTs of length over . We decompose the DFT into “short” DFTs of length as follows. Let and , and write where for and . We use the algorithm , where for we take to be the algorithm based on Bluestein's method (discussed immediately before (9.4)), and where is the usual Cooley–Tukey algorithm over . Let be the cost of a single invocation of (or of the corresponding inverse transform ). By (2.4) we have
The cost of precomputing the necessary root tables is only . By definition . From (9.4) and the estimate , the first term becomes
The contribution to from all terms involving is
so
Denoting by the cost of a bivariate DFT of length over , we thus have (ignoring the transposition costs, which were included earlier)
Moreover, since
we get
We must perform bivariate DFTs; the bound (9.5) then follows exactly as in the proof of Theorem 6.4.
For large , we have , so . Thus there exists a constant such that for large , and we may take .
Proof of Theorem 1.2. Follows from Theorem 9.6 and Proposition 5.3, analogously to the proof of Theorem 1.1.
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