1. Introduction
Due to their simple algebraic structures, cyclic codes over finite fields constitute one of the classes of linear codes that have attracted considerable attention in the literature. This has inspired researchers to explore more general classes of linear codes that have a similar structure. Generalizing the shift constant of cyclic codes introduces the class of constacyclic codes [
1], and generalizing the shift index of cyclic codes introduces the class of quasi-cyclic (QC) codes [
2]. Eventually, the class of quasi-twisted (QT) codes appeared to generalize QC and constacyclic codes. A
-QT code of index
ℓ is a linear code that is invariant under the
-constacyclic shift of
ℓ coordinates. In particular, a one-QT code is QC, while a
-QT code of index
is
-constacyclic.
Having a code with the largest minimum distance is beneficial for real-world communication systems, as the code’s error-correction capability is proportional to its minimum distance. Databases such as [
3] list codes with the best-known parameters. However, as pointed out by ref. [
4], determining the minimum distance of a linear code is an NP-hard problem. As such, researchers have focused on providing lower and upper bounds to the minimum distance as an alternative. The Hartmann–Tzeng bound and the Bose–Chaudhuri–Hocquenghem (BCH) bound are two of the well-known lower bounds for cyclic codes. The BCH bound for cyclic codes was extended to constacyclic codes in [
5,
6]. Furthermore, Lally [
7] studied a lower bound on the minimum distance of QC codes. Specifically, a cyclic code over some extension field was built, and its minimum distance was employed to determine a lower bound on the minimum distance of the given QC code. Therefore, to compute this bound, it is necessary to perform calculations in an extension field of the field over which the QC code is defined. Lower bounds for single-generator QT codes are presented in [
5,
8], where the code generator has a specific pattern. These bounds can also be calculated by employing the BCH bound for constacyclic codes.
This paper begins with a proof of the BCH bound for constacyclic codes. We prove this bound for completeness, as all subsequent bounds depend on it. Our study focuses on the lower bounds on the minimum distance of single-generator QT codes. We establish an analogous bound to Lally’s result [
7] for the QT code. Specifically, we associate a constacyclic code over an extension field to each QT code, and then we employ the BCH bound of constacyclic codes to determine a minimum-distance lower bound for the QT code. For general single-generator QT codes, some weaknesses, however, are notable for this bound. We propose an alternative lower bound to tackle this weakness. To this end, we implicitly utilize the Chinese remainder theorem (CRT) to decompose the QT code into a direct sum of minimal QT subcodes. The proposed bound employs this decomposition in conjunction with the BCH bound for constacyclic codes. We outline two advantages for the proposed bound: firstly, it does not presuppose any specific form for the code generator; secondly, all calculations are carried over the same field of the code alphabet, not an extension field.
We compare our proposed lower bound with those presented in [
5,
7,
8] by analyzing several examples with various parameters. Although the lower bound proposed in [
5] assumes a specific form for the code generator, we do not assume any specific form in our bound. We show that when the code generator has the designed form in [
5], the two bounds coincide. This shows that the proposed bound generalizes that in [
5]. Theorem III.2 in [
8] presents another lower bound on the minimum distance of any single-generator QT code; however, we contradict this bound with Example 6 below.
The remainder of this paper is divided into the following sections: In
Section 2, the algebraic structures of constacyclic and QT codes are reviewed.
Section 3 presents the lower bound of single-generator QT codes, which mimics the work in [
7] on QC codes.
Section 4 provides the proposed bound in detail.
Section 5 then examines the proposed bound on several numerical examples. Finally, the paper is concluded in
Section 6.
2. Algebraic Structures of QT Codes
A
cyclic code of length
m over a finite field
is a linear subspace of
that is invariant under the
shift transformation
Cyclic codes form a subclass of constacyclic codes. For a nonzero
, a
-constacyclic code over
of length
m is a linear subspace of
that is invariant under the
constacyclic shiftThe polynomial representation to the codewords of constacyclic codes endows them with the algebraic structure of an ideal of a ring. Let be the ring of all polynomials over of degrees less than m, where the addition and multiplication are defined as modulo . The codeword of a -constacyclic code is mapped to the polynomial in . Hence, is regarded as an ideal of . Conversely, any ideal of corresponds to a constacyclic code. Throughout this paper, we refer to the -constacyclic code over of length m as an ideal of . We call the shift constant of . It follows that a cyclic code is constacyclic with a shift constant equal to unity.
By increasing the number of coordinates that must be shifted for the code to be invariant under this shifting, cyclic codes can be generalized to
quasi-cyclic (QC) codes. A QC code over
of length
n is a linear subspace of
that is invariant under the cyclic shift by
ℓ coordinates. The length of a QC code is divisible by
ℓ, that is,
, where
ℓ is the
index, and
m is the
co-index. A typical QC codeword can be subdivided as
and represented as a polynomial vector
, where
for
. In the polynomial representation, a QC code is an
submodule of
. In turn, each
submodule of
corresponds to a QC code over the
of index
ℓ and co-index
m. It follows that a cyclic code is a QC of index
.
Constacyclic codes are not a subclass of QC codes. Therefore, it is convenient to introduce the class of quasi-twisted (QT) codes, which generalizes both constacyclic and QC codes. For a nonzero , a -QT code over of length , index ℓ, and co-index m is a linear subspace of that is invariant under the -constacyclic shift by ℓ coordinates. Again, we refer to as the shift constant of . QT codes, as with QC codes, have a polynomial representation, which turns them into -submodules of , where . Throughout this paper, we regard QT codes as submodules of . It follows that a constacyclic code is a QT of index , but a QC code is 1-QT. A QT code is said to be a single generator if there exists such that . For each , is an element of . Nevertheless, in the sequel, we may mean that is the equivalent polynomial of degree less than m in the polynomial ring . This happens throughout the paper, and the reader should be able to identify to which polynomial we are referring, whether it belongs to or . For instance, to determine the greatest common divisor of some polynomials, denoted , or the roots of a polynomial, it is beneficial to think of these polynomials as elements of the principal ideal domain .
We conclude this section by summarizing some of the results in [
7]. Suppose
is a single-generator QC code over
of length
generated by
. In [
7],
is associated with a cyclic code
over
of length
m as follows: Fix an element
of degree
ℓ over
. By viewing
as a polynomial in
, define
, which is now considered an element of
. The ideal generated by
in
is the cyclic code
associated with
. We can write
; therefore,
is an
submodule of
. On the other hand,
can be thought of as an
module. The codewords of
are in one-to-one correspondence with the elements of the
submodule of
generated by
. This correspondence given by
. It is shown in [
7] that the dimension of
as an
-vector space is
Furthermore, a lower bound on the minimum Hamming distance
of
is given by
where
is the linear code over
of length
ℓ generated by the vector equivalent of the coefficients of
with respect to the
basis
of
. In the next section, we generalize this result from QC codes to QT codes.
3. Constacyclic and QT Codes Bound
The objective of this section is to provide the expected generalization of the result in [
7] for single-generator QT codes. It seems reasonable that the equivalent bound of QT codes would require a BCH bound for constacyclic codes, because the bound of QC codes requires a BCH bound for cyclic codes. In fact, the BCH bound for constacyclic codes is essential to all lower bounds that we present in this paper as well as for use in our generalization of [
7]. As a result, we begin this section with a simple proof of this BCH bound for the convenience of the reader.
Let be a -constacyclic code over of length m, which is regarded as an ideal in the quotient ring . Because is a principal ideal ring, there is a generator polynomial , such that . Instead of introducing some notations to differentiate between polynomials in and their images in , we leave it up to the reader to determine our meaning. Hence, the generator polynomial of as an element of , which we also denote by , is defined as the unique monic codeword of of a minimum degree. It is noted that divides . Suppose that and are -constacyclic codes with generator polynomials and , respectively. Accordingly, is a -constacyclic code with the generator polynomial , which is the greatest common divisor of and in . In general, let , , be a collection of -constacyclic codes indexed by a set , and let be the generator polynomial of for each . Then, is the generator polynomial of .
Lemma 1. Let be a collection of elements of indexed by a set . Then, is the smallest λ-constacyclic code that contains , where Proof. Let be any -constacyclic code that contains . Then for each . Consequently, . On the other hand, is a -constacyclic code. As a result, it is the smallest code that contains . □
Recall that the BCH bound on the minimum distance of a cyclic code of length
m necessitates the determination of the
mth roots of unity in some extension field of
. Similarly, finding the
m zeros of
in an extension field
of
is required to determine the BCH bound on the minimum distance of a
-constacyclic code of length
m. Specifically, assume
m and
q are coprime, the multiplicative order of
is
r, and the multiplicative order of
q modulo
is
e. Then, the splitting field of
is
. The
m zeros of
are precisely
, where
and
are the
mth roots of
and unity, respectively, in
. In other words, if
is a primitive element of
and
for some integer
l, then the zeros of
are
Definition 1. Suppose m and q are coprime. Let denote the generator polynomial of a λ-constacyclic code over of length m. Furthermore, let β and γ be the mth roots of λ and unity, respectively, in . Define , where δ is the largest positive integer such thatfor some integer . If for all , define . It is important to select the integer
b of Definition 1 with care in order to maximize the number of consecutive
mth roots of
that are zeros of
, as this can lead to a better lower bound on the minimum distance; this is shown in the following theorem. The following theorem presents a BCH-like bound for constacyclic codes. The proof, utilizing elementary linear algebra techniques, can be found in [
6].
Theorem 1. Suppose m and q are coprime. Let be a λ-constacyclic code over of length m with generator polynomial . Then, the minimum distance of is at least .
We are now ready to imitate Lally’s work [
7] in order to give a lower bound for QT codes. Let
be a single-generator
-QT code over
of length
and index
ℓ. Suppose that
, where
is a generator of
. Fix an element
of degree
ℓ over
; hence,
is a basis of
as a vector space over
. Define
as an element of
. This polynomial generates an ideal of
that is a
-constacyclic code over
of length
m, which we denote by
. Each codeword
is linked with the codeword
. A lower bound on the minimum distance of
can be found with Theorem 1, which may be utilized to establish a lower bound on the minimum distance of
(cf. Theorem 3 in [
7]). The result is precisely as follows:
Theorem 2. Let be a λ-QT code over of index ℓ and co-index m, generated by . Let ; and let denote the λ-constacyclic code over of length m generated by . Let be the linear code over of length ℓ generated by the vectors equivalent to the coefficients with respect to the basis of . Then, Proof. Assume that is an arbitrary nonzero codeword, where and . Then, . This can be rearranged as , where . Because the latter is a codeword of , the number of nonzero for is at least . There is also a polynomial , such that . This means that each is an -linear combination of . Therefore, for every . The result follows, because for every nonzero , we have . □
The dimension of the QT code described in Theorem 2 can be determined using the following theorem, the proof of which is nearly identical to that of Theorem 2 in [
7].
Theorem 3. Suppose is a λ-QT code over of index ℓ and co-index m generated by . The dimension of as an -vector space is Example 1. Consider the single-generator 2-QT code over of index , co-index , and generator In , . Theorem 3 states that the dimension of is 6. Let α denote a zero of the primitive polynomial . The 2-constacyclic code described in Theorem 2 is generated by . We employ Theorem 1 to determine a lower bound to . In fact, the splitting field of is . If ζ is a zero of the primitive polynomial , then . We set and , which are the 7th root of 2 and unity, respectively. Observe that are the zeros of . By Definition 1, . It follows from Theorem 1 that . The generator matrix of the code over of length described in Theorem 2 is We find that ; hence, . The exact minimum distance of is found to be .
In Example 1, the polynomial has a degree of 4, and all of its zeros are zeros of . Thus, divides . In this situation, we observe that Theorem 2 yields an acceptable bound. However, when does not divide , we observe that the lower bound provided by Theorem 2 is a weak bound in most cases. We show this with the following example, which employs a randomly generated QT code.
Example 2. We examine a single-generator ω-QT code over of index , co-index , where ω is a zero of . We randomly construct which is described by the starting values and the linear recurrence relation for . This results inwhose polynomial representation is as follows Let α denote a zero of . The ω-constacyclic code described in Theorem 2 is generated by = It is only necessary to observe that is coprime to . Then, ; hence, . The generator matrix of code over of length described in Theorem 2 is We find that ; hence, . The exact minimum distance of is found to be .
Example 2 explores one of the weaknesses of Theorem 2. The other disadvantages of using Theorem 2 are that it requires calculations over an extension field as well as determining the exact minimum distance of , which is generally NP-hard. In response to these weaknesses, we propose a novel lower bound on the minimum distance of single-generator QT codes in the next section.
4. Novel Bound to QT Codes
We expect the lower bound given in Theorem 2 to be far from the exact minimum distance for a general single-generator QT code. This is demonstrated in Example 2, especially when the polynomial
, as an element of
, is coprime to
. In effect, this results in
and, hence,
. In this section, we deal with this problem by establishing a completely new lower bound for single-generator QT codes. We apply the CRT to decompose any single-generator QT code before employing Theorem 1 at each polynomial coordinate. Throughout this section,
denotes a single-generator QT code over
of index
ℓ and co-index
m, where
m and
q are coprime. Suppose
is a generator of
, where
. Then,
, which is the cyclic
submodule of
generated by
. Because
m and
q are coprime,
, where
are distinct irreducible polynomials in
. Now, we define
for each
, where
Clearly,
is a single-generator QT code annihilated by
. Then,
is a minimal QT subcode of
because
. To exclude constituents
, which are not involved in the construction of
, we define the set
, or, equivalently, using (
1), we have the following alternative definition:
The following result shows that can be decomposed to the direct sum of the minimal QT subcodes through employing the CRT implicitly.
Lemma 2. With the above notation, .
Proof. Because
for every
,
. Because
in
, there are polynomials
such that
. Consequently,
Accordingly,
; therefore,
. Now suppose that
, then
Then, annihilates . However, is annihilated by ; then, . For every arbitrary , = . Because , we infer that , so . Therefore, for every , and we conclude that . □
Recall that the power set of X is the set of all subsets of X. We denote the power set of X after excluding the empty set by . That is, is the set of all nonempty subsets of X. We now prove that every nonzero codeword of corresponds to a unique element of .
Lemma 3. For each nonzero , there exist a unique and nonzero polynomials (for ) such that .
Proof. From Lemma 2,
. Thus, we have
for some
. The result is achieved by defining
□
The following definition attempts to avoid the zero coordinates for each
. This is achieved by defining the set
where
denotes the
jth entry of
, or, equivalently, using (
1), we have the following alternative definition:
Definition 2. For each , define In other words, if is the associated element in with the codeword , as shown by Lemma 3, then keeps a record of the nonzero coordinates of . The following result gives a lower bound on the weight of any codeword .
Lemma 4. Assume is a nonzero codeword of , and let , where and for each . Then, Proof.
Now, for each
, let
be the smallest
-constacyclic code over
of length
m that contains
. From Lemma 1,
generates
. We then deduce from Theorem 1 that
. Therefore,
because
. The result follows from (
3). □
Lemma 3 asserts that each codeword of corresponds to an element . However, Lemma 4 provides a lower bound for each codeword from its corresponding . Hence, taking into account all elements of ensures that no codeword is missed. Therefore, a lower bound on is stated as follows:
Theorem 4. Let be a single-generator QT code over of index ℓ and co-index m, where m and q are coprime. With the notation introduced above,where τ runs over . Proof. The following lower bound on
is now immediately available from the previous discussion:
Therefore, (
4) follows from (
1) by observing that
This is because if and only if . □
Although Theorem 4 is applied to several numerical examples in
Section 5, we conclude this section by applying it to Example 1. Recall that Theorem 2 provides a lower bound of 8 on the minimum distance of the code presented in Example 1. Indeed, as mentioned at the beginning of this section, Theorem 2 provides an acceptable lower bound when
divides
, which is the case in Example 1. However, to establish a fair comparison between Theorems 2 and 4, they must be examined on a randomly generated code, which is performed in the next section.
Example 3. We proceed with the QT code considered in Example 1. Recall thatgenerates . The irreducible factors of in are and . Equation (2) implies that ; hence, . By Definition 2, . By Theorem 4, we findwhere by Definition 1. In Examples 1 and 3, recall that the polynomial defined in Theorem 2 divides . Although these examples show that Theorems 2 and 4 give the same lower bound, Theorem 4 has the advantage in that it does not require any calculations over any extension fields nor does it require calculating the minimum distance of another linear code: . On the other hand, a disadvantage of the lower bound of Theorem 4 appears when decomposes to many irreducible factors. In fact, the size of exponentially increases with the size of X.
5. Numerical Examples
We have four goals in this section: We first compare the proposed bound to the one provided in Theorem 2. Indeed, as demonstrated by Example 2, the latter bound has some weaknesses. Therefore, we inspect the bound of Theorem 4 on the code of Example 2. Second, we investigate the lower bound suggested in Theorem 4 for codes with different indices. Third, we consider the lower bound introduced in Theorem 3.2 of [
5]. In fact, this bound does not apply to all single-generator QT codes: it only suits a specific form for the code generator. We consider this specific form and prove that Theorem 4 and Theorem 3.2 in [
5] are equivalent under the assumption of this form. Therefore, we may argue that the proposed bound generalizes that in Theorem 3.2 in [
5] since it does not require any specific form for the code generator. Lastly, we present a counterexample to the lower bound provided in Theorem III.2 in [
8].
We begin with the following example to examine Theorem 4 on a nonbinary QT code. Remember that the QT code introduced in Example 2 shows a flaw in the lower bound given by Theorem 2. Specifically, Theorem 2 gives a lower bound of 2 to a QT of minimum distance 17. We need to determine what lower bound Theorem 4 can achieve for the same code.
Example 4. Let be the QT code of Example 2 generated by the randomly chosen codeword Because and , , , . Let ζ be a zero of the primitive polynomial . Then, , while and are the primitive 5th root of unity and ω, respectively. The zero of is , the zeros of are , and the zeros of are .
From Theorem 3, the dimension of is . Equation (2) implies that .
- 1.
For , by Definition 2, and .
- 2.
For , by Definition 2, and .
- 3.
For , by Definition 2, and .
- 4.
For , by Definition 2, and .
- 5.
For , by Definition 2, and .
- 6.
For , by Definition 2, and .
- 7.
For , by Definition 2, and .
Therefore, , where the minimum distance of is found to be .
We now examine Theorem 4 on several binary QC codes of the same length with varied indices and co-indices.
Example 5. We randomly generate the binary codeword
of length 45. This codeword is interpreted as a generator of several QC codes of length 45 and different indices. We determine the lower bound on provided by Theorem 4 for different ℓ values. - 1.
Let . Then, , , and . The zeros of are , and the zero of is . In polynomial representation, From Theorem 3, the dimension of is . Equation (2) implies that . - (a)
For , .
- (b)
For , .
- (c)
For , .
Therefore, . In fact, we find .
- 2.
Let . Then , , and . Let ζ be a zero of the primitive polynomial . Then, is the primitive 5th root of unity. The zeros of are , and the zero of is . In polynomial representation, From Theorem 3, the dimension of is . Equation (2) implies that . - (a)
For , .
- (b)
For , .
- (c)
For , .
Therefore, . In fact, we find .
- 3.
Let . Then, , , , and . Let ζ be a zero of the primitive polynomial . Then, is the primitive 9th root of unity. The zeros of are , the zeros of are , and the zero of is . In polynomial representation, From Theorem 3, the dimension of is . Equation (2) implies that . For or , the right side of (4) has the smallest value. Therefore, . We found . - 4.
Let . Then , , , , , and . Let ζ be a zero of the primitive polynomial . The zeros of are , the zeros of are , the zeros of are , the zeros of are , and the zero of is . In polynomial representation, From Theorem 3, the dimension of is . Equation (2) implies that . For , the right side of (4) has the smallest value. Therefore, . We found .
Theorem 3.2 in [
5] introduces a lower bound on the minimum distance of a single-generator QT code
, where the code generator is assumed to be of the form
divides
, and
is coprime to
for
, as elements of
. It was shown that
under this particular form (
6) of
. In fact, this lower bound is limited to single-generator QT codes with this generator form. For instance, it does not apply to the code in Example 4. Specifically, (
5) forces us to choose
, so
is not coprime to
. This reinforces our proposed bound, which has the advantage of being applicable to any single-generator QT code because Theorem 4 makes no assumptions about the code generator. Furthermore, we prove in Corollary 1 below that Theorem 4 generalizes Theorem 3.2 in [
5]. To this end, assuming the code has generator (
6), we show that the lower bound of Theorem 4 is reduced to that of Theorem 3.2 in [
5].
Corollary 1. Let be a QT code over of index ℓ and co-index m, where m and q are coprime. Assume withwhere divides , and is coprime to for . Then Proof. Suppose that
and, without loss of generality, that
for some
. Then,
for
and
for
. Hence,
. The condition on
implies that for every
,
. Consequently, the value of the right side of (
4) is the smallest when
. From Theorem 4, we find
□
We conclude this section by contradicting Theorem III.2 in [
8], which provides a lower bound on the minimum distance of any single-generator QT code. Specifically, Theorem III.2 in [
8] states that
for any single-generator QT code with generator
. This theorem is contradicted by the following example.
Example 6. Let be the binary QC code of index and co-index generated by From Theorem III.2 in [8], . However, the codeword has a weight of 4, which contradicts the lower bound. On the other hand, Theorem 4 demonstrates that . This can be shown as follows: We have , , and . Equation (2) implies that . - 1.
For , and .
- 2.
For , and .
- 3.
For , and .
Therefore, ; hence, .