# Quantum Codes from Linear Codes over Finite Chain Rings

Xiusheng Liu
School of Mathematics and Physics,
Hubei Polytechnic University
Huangshi, Hubei 435003, China,
Email<EMAIL_ADDRESS>
Hualu Liu
School of Mathematics and Physics,
Hubei Polytechnic University
Huangshi, Hubei 435003, China,
Email<EMAIL_ADDRESS>

###### Abstract

In this paper, we provide two methods of constructing quantum codes from
linear codes over finite chain rings. The first one is derived from the
Calderbank-Shor-Steane (CSS) construction applied to self-dual codes over
finite chain rings. The second construction is derived from the CSS
construction applied to Gray images of the linear codes over finite chain ring
$\mathbb{F}_{p^{2m}}+u\mathbb{F}_{p^{2m}}$. The good parameters of quantum
codes from cyclic codes over finite chain rings are obtained.

Keywords Quantum codes $\cdot$ Cyclic codes $\cdot$ Dual-containing codes,

##  1 Introduction

Quantum codes were introduced to protect quantum information from decoherence
and quantum noise. A main obstacle to complete quantum communication is
decoherence of quantum bits caused by inevitable interaction with
environments. Quantum codes provide an efficient way to overcome decoherence.
After the works of Shor [1] and Steane [2], the theory of quantum codes has
been progressed rapidly. Calderbank et al. [3] provided a systematic
mathematical scheme to construct quantum codes from classical Hermitian dual-
containing codes over finite field $\mathbb{F}_{4}$.

Some authors constructed quantum codes by using linear codes over finite
rings. For example, in [4], Qian et al. gave a new method to obtain self-
orthogonal codes over $\mathbb{F}_{2}$. They gave a construction for quantum
codes starting from cyclic codes over finite ring,
$\mathbb{F}_{2}+u\mathbb{F}_{2},u^{2}=0$. In [5], Ashraf and Mohammad
constructed quantum codes from cyclic codes over
$\mathbb{F}_{3}+v\mathbb{F}_{3}$. In this paper, we continue to study quantum
codes which are derived from finite chain ring.

Hereafter, $p$ is a prime. The purpose of this paper is to consider liner
codes over finite chain rings to obtain good quantum codes. In Section 2, we
review some concepts and properties about quantum codes over finite fields. In
section 3, we first give the necessary background materials on finite chain
rings. Then a construction for quantum codes from self-dual codes over finite
chain rings is given. In the final section, for special finite chain ring
$R=\mathbb{F}_{p^{2m}}+u\mathbb{F}_{p^{2m}}$, we define a new Gray map $\Phi$
from $R^{n}$ to $\mathbb{F}_{p^{2m}}^{2n}$, Gray weights of elements of
$R^{n}$, Gray distance $d_{G}(C)$ and Hermitian dual $C^{\perp_{H}}$ with
respect to Gray weight and the Hermitian inner product in the linear code $C$
over $R$. We prove that $\Phi(C^{\perp_{H}})=\Phi(C)^{\perp_{H}}$, and give a
method to derive Hermitian dual-containing codes over $\mathbb{F}_{p^{2m}}$ as
Gray images of linear codes over $\mathbb{F}_{p^{2m}}+u\mathbb{F}_{p^{2m}}$.
The parameters of quantum codes are obtained from cyclic codes over $R$.

##  2 Review of Symmetric Quantum Codes

In this section, we recall some basic concepts and results of symmetric
quantum codes, necessary for the development of this work. For more details,
we refer to [6,7].

Let $\mathbb{F}_{q}$ be the finite field with $q=p^{2m}$, where $p$ is a prime
number and $m\geq 1$ is an integer. Let $V_{n}$ be the Hilbert space
$V_{n}=\mathbb{C}^{q^{n}}=\mathbb{C}^{q}\otimes\cdots\otimes\mathbb{C}^{q}$.
Let $\mid x\rangle$ be the vectors of an orthonormal basis of
$\mathbb{C}^{q^{n}}$, where the labels $x$ are elements of $\mathbb{F}_{q}$.
Then $V_{n}$ has the following orthonormal basis
$\\{|c\rangle=|c_{1}c_{2}\cdots
c_{n}\rangle=|c_{1}\rangle\otimes|c_{2}\rangle\otimes\cdots\otimes|c_{n}\rangle:~{}c=(c_{1},c_{2},\ldots,c_{n})\in\mathbb{F}_{q}^{n}\\}$.

Consider $a,b\in\mathbb{F}_{q}$, the unitary linear operators $X(a)$ and
$Z(b)$ in $\mathbb{C}^{q}$ are defined by $X(a)|x\rangle=|x+a\rangle$ and
$Z(b)|x\rangle=w^{tr(bx)}|x\rangle$, respectively, where $w=exp(2\pi i/p)$ is
a primitive $p$-th root of unity and tr is the trace map from $\mathbb{F}_{q}$
to $\mathbb{F}_{p}$.

Let $\mathbf{a}=(a_{1},\ldots,a_{n})\in\mathbb{F}_{q}^{n}$, we write
$X(\mathbf{a})=X(a_{1})\otimes\cdots\otimes X(a_{n})$ and
$Z(\mathbf{a})=Z(a_{1})\otimes\cdots\otimes Z(a_{n})$ for the tensor products
of $n$ error operators. The set
$E_{n}=\\{X(\mathbf{a})Z(\mathbf{b}):\mathbf{a},\mathbf{b}\in\mathbb{F}_{q}^{n}\\}$
is an error basis on the complex vector space $\mathbb{C}^{q^{n}}$ and we set
$G_{n}=\\{w^{c}X(\mathbf{a})Z(\mathbf{b}):\mathbf{a},\mathbf{b}\in\mathbb{F}_{q}^{n},c\in\mathbb{F}_{p}\\}$
is the error group associated with $E_{n}$.

######  Definition 2.1.

A $q$-ary quantum code of length $n$ is a subspace $Q$ of $V_{n}$ with
dimension $K>1$. A quantum code $Q$ of dimension $K>2$ is called symmetric
quantum code (SQC) with parameters $((n,K,d))_{q}$ or $[[n,k,d]]_{q}$,where
$k=log_{q}K$ if $Q$ detect $d-1$ quantum digits of errors for $d\geq 1$.
Namely, if for every orthogonal pair $|u\rangle,|v\rangle$ in $Q$ with
$<u|v>=0$ and every $e\in G_{n}$ with $W_{Q}(e)\leq d-1$, $|u\rangle$ and
$e|v\rangle$ are orthogonal, i.e.,$<u|e|v>=0$. Such a quantum code is called
pure if $<u|e|v>=0$ for any $|u\rangle$ and $|v\rangle$ in $Q$ and any $e\in
G_{n}$ with $1\leq W_{Q}(e)\leq d-1$. A quantum code $Q$ with $K=1$ is always
pure.

Let us recall the SQC construction:

######  Theorem 2.2.

$[7,\mathrm{Lemma}~{}20]$ Let $C_{i}$ be a classical linear code with
parameters $[n,k_{i},d_{i}]_{q}$ and $C_{i}^{\perp}\subseteq
C_{1+(i~{}mod~{}2)}$$(i=1,2)$. Then there exists an SQC Q with parameters
$[[n,k_{1}+k_{2}-n,\geq d]]_{q}$ that is pure to $min\\{d_{1},d_{2}\\}$, where
$d=min\\{wt(c):c\in(C_{1}\backslash C_{2}^{\perp})\cup(C_{2}\backslash
C_{1}^{\perp})\\}$.

######  Corollary 2.3.

If $C$ be a classical linear $[n,k,d]_{q}$ code containing its dual
$C^{\perp}\subseteq C$, then there exists an SQC Q with parameters
$[[n,2k-n,\geq d]]_{q}$ that is pure to $d$.

To see that an SQC is good in terms of its parameters, we have to introduce
the quantum Singleton bound ( See [6]).

######  Theorem 2.4.

Let $C$ be an SQC with parameters $[[n,k,d]]_{q}$. Then $k\leq n-2d+2$.

If an SQC $Q$ with parameters $[[n,k,d]]_{q}$ attains the quantum Singleton
bound $k=n-2d+2$, then it is called an SQC maximum-distance-separable (SQCMDS)
code.

######  Definition 2.5.

An SQC $Q$ with parameters $[[n,k,d]]_{q}$ is called a near quantum maximum
distance separable (SQCNMDS) code if it satisfies $2d\geq n-k$.

######  Corollary 2.6.

Let $C$ be an $[n,k,d]_{q}$ classical code containing its dual,
$C^{\perp}\subseteq C$. Then

  * (1)

$k\geq\lceil\frac{n}{2}\rceil.$

  * (2)

If $C$ is an MDS code, then there exists an $[[n,2k-n,d]]_{q}$ SQCMDS code.

Proof. (1) Since $C$ is a $k$-dimensional subspace of $\mathbb{F}_{q}^{n}$,
$C^{\perp}$ is a $(n-k)$-dimensional subspace of $\mathbb{F}_{q}^{n}$. It
follows that $n-k\leq k$ by $C^{\perp}\subseteq C$. Therefore,
$k\geq\lceil\frac{n}{2}\rceil.$

(2) If $C$ is an $[n,k,d]_{q}$ classical MDS codes containing its
dual,$C^{\perp}\subseteq C$, then Corollary 2.3 implies the existence of a
quantum $[[n,2k-n,\geq d]]_{q}$ code $Q$. Theorem 2.4 shows that the minimum
distance of $Q$ is $\leq\frac{n-(2k-n)+2}{2}=n-k+1$, so $Q$ is an
$[[n,2k-n,d]]_{q}$ SQCMDS code.

######  Corollary 2.7.

Let $C$ be an $[n,k,d]_{q}$ classical code containing its dual,
$C^{\perp}\subseteq C$, and $2d\geq n-k$. Then there exists an $[[n,2k-n,\geq
d]]_{q}$ SQCNMDS code.

Proof. If $C$ is an $[n,k,d]_{q}$ classical codes containing its
dual,$C^{\perp}\subseteq C$, then Corollary 2.3 implies the existence of a
quantum $[[n,2k-n,d_{1}]]_{q}$ code $Q$, where $d_{1}\geq d$. So $Q$ is an
$[[n,2k-n,\geq d]]_{q}$ SQCNMDS code.

##  3 SQC from to the Linear Codes over Finite Chain Rings

Constructions of quantum codes are exhaustively investigated in the
literature. As mentioned in Section 1, some authors have exhibited families of
optimal codes. However, many of these techniques are based on the construction
of classical codes over finite fields.

In this section, we use self-dual codes over finite chain rings to construct
SQC.

We begin with some definitions and properties about finite chain rings
(see[8,9]). Let $R$ be a finite commutative ring with identity. A nonempty
subset $C\subseteq R^{n}$ is called a linear code of length $n$ over $R$ if it
is an $R$-submodule of $R^{n}$. Throughout this section, all codes are assumed
to be linear.

A commutative ring is called a chain ring if the lattice of all its ideals is
a chain. It is well known that if $R$ is a finite chain ring, then $R$ is a
principal ideal ring and has a unique maximal ideal
$\langle\gamma\rangle=R\gamma=\\{r\gamma\,|\,r\in R\\}$. Its chain of ideals
is

$R=\langle\gamma^{0}\rangle\supset\langle\gamma^{1}\rangle\supset\cdots\langle\gamma^{t-1}\rangle\supset\langle\gamma^{t}\rangle=\\{0\\}.$

The integer $t$ is called the nilpotency index of $\gamma$. Note that the
quotient $R/{\langle\gamma\rangle}$ is a finite field $\mathbb{F}_{q}$, where
$q$ is a power of a prime $p$. There is a natural homomorphism from $R$ onto
$\mathbb{F}_{q}=R/\langle\gamma\rangle$, i.e.,

${}^{-}:R\longrightarrow\mathbb{F}_{q}=R/\langle\gamma\rangle,~{}~{}~{}r\mapsto
r+\langle\gamma\rangle=\overline{r},~{}~{}~{}~{}\hbox{for any $r\in R$}.$

We need the following lemma (see [9]).

######  Lemma 3.1.

Let $R$ be a finite chain ring with maximal ideal $\langle\gamma\rangle$. Let
$V\subset R$ be a set of representatives for the equivalence classes of $R$
under congruence modulo $\gamma$. Then

  * (1)

for any $v\in R$ there exist unique $v_{0},\ldots,v_{t-1}\in V$ such that
$v=\sum_{i=0}^{t-1}v_{i}\gamma^{i}.$

  * (2)

$|V|=|R/\gamma|=|\mathbb{F}_{q}|$.

The natural homomorphism from $R$ onto $\mathbb{F}_{q}=R/\langle\gamma\rangle$
can be extended naturally to a projection from $R^{n}$ onto
$\mathbb{F}_{q}^{n}$. For an element $c\in R^{n}$, let $\overline{c}$ be its
image under this projection. Given $r\in R$ and $c\in R^{n}$, we denote by
$rc$ the usual multiplication of a vector by a scalar. Let $C$ be a code of
length $n$ over $R$. We define $\overline{C}=\big{\\{}\overline{c}\,|\,c\in
C\big{\\}}$ and $(C:r)=\big{\\{}e\in R^{n}\,|\,re\in C\big{\\}}$, where $r$ is
an element of $R$.

######  Definition 3.2.

To any code $C$ over $R$, we associate the tower of codes

$C=(C:\gamma^{0})\subseteq(C:\gamma)\subseteq\ldots\subseteq(C:\gamma^{t-1})$

over $R$ and its projection to $\mathbb{F}_{q}$,

$\overline{C}=\overline{(C:\gamma^{0})}\subseteq\overline{(C:\gamma)}\subseteq\ldots\subseteq\overline{(C:\gamma^{t-1})}.$

The following definitions and results can be found in [9].

######  Definition 3.3.

Let $C$ be a linear code over $R$. A generator matrix $G$ for $C$ is said to
be in standard form if after a suitable permutation of the coordinates,

$G=\begin{pmatrix}I_{k_{0}}&A_{0,1}&A_{0,2}&A_{0,3}&\cdots&A_{0,t-1}&A_{0,t}\\\
0&\gamma I_{k_{1}}&\gamma A_{1,2}&\gamma A_{1,3}&\cdots&\gamma
A_{1,t-1}&\gamma A_{1,t}\\\
0&0&\gamma^{2}I_{k_{2}}&\gamma^{2}A_{2,3}&\cdots&\gamma^{2}A_{2,t-1}&\gamma^{2}A_{2,t}\\\
\cdots&\cdots&\cdots&\cdots&\cdots&\cdots&\cdots\\\
0&0&0&0&\cdots&\gamma^{t-1}I_{k_{t-1}}&\gamma^{t-1}A_{t-1,t}\\\
\end{pmatrix}=\begin{pmatrix}A_{0}\\\ \gamma A_{1}\\\ \vdots\\\
\gamma^{t-1}A_{t-1}\end{pmatrix}.~{}~{}~{}~{}~{}~{}~{}(3.1)$

We associate to $G$ the matrix

$~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}A=\begin{pmatrix}A_{0}\\\
A_{1}\\\ \vdots\\\
A_{t-1}\end{pmatrix}.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(3.2)$

For a code $C$, we define the rank of $C$, denoted by rank$(C)$, to be the
minimum number of generators of $C$ and the free rank of $C$, denoted by free
rank$(C)$, to be the maximum of the ranks of free $R$-submodules of $C$. Codes
where the rank is equal to the free rank are called free codes.

Let $C$ be a linear code over $R$. We denote by $k(C)$ the number of rows of a
generating matrix $G$ in standard form for $C$, and for $i=1,2,\ldots,t-1$ we
denote by $k_{i}(C)$ the number of rows of $G$ that are divisible by
$\gamma^{i}$ but not by $\gamma^{i+1}$.

Clearly, rank$C=k(C)=\sum_{i=0}^{t-1}k_{i}(C).$

It is well known (see [10]) that for codes $C$ of length $n$ over any alphabet
of size $m$

$d_{H}(C)\leq n-log_{m}(|C|)+1.$

Codes meeting this bound are called MDS (Maximum Distance Separable) codes.

Further if C is linear, then

$d_{H}(C)\leq n-rank(C)+1.$

Codes meeting this bound are called MDR (Maximum Distance with respect to
Rank) codes.

######  Lemma 3.4.

Let $C$ be a linear code over $R$ with a generator matrix $G$ in standard form
and let $A$ be as in $(3.2)$.

$(1)$For $0\leq i\leq t-1$, $\overline{(C:\gamma^{i})}$ has generator matrix

$\overline{G_{i}}=\left(\begin{array}[]{c}\overline{A_{0}}\\\ \vdots\\\
\overline{A_{i}}\end{array}\right)$

and $\mathrm{dim}\overline{(C:\gamma^{i})}=k_{0}(C)+\cdots+k_{i}(C)$.

$(2)$ For $0\leq i\leq t-1$,
$\overline{(C^{\perp}:\gamma^{i})}=\overline{(C:\gamma^{t-1-i})}^{\perp}$. We
have $k(C^{\perp})=n-k_{0}(C),k_{0}(C^{\perp})=n-k(C)$, and
$k_{i}(C^{\perp})=k_{t-i}(C)$, for $i=1,\ldots,t-1$.

$(3)$ $d_{H}(C)=d_{H}\overline{(C^{\perp}:\gamma^{t-1})}$.

$(4)$ If $C$ is an MDR code over $R$, then
$\overline{(C^{\perp}:\gamma^{t-1})}$ is an MDS code over
$\mathbb{F}_{q}=R/\langle\gamma\rangle$.

We have an important observation that proves to be rather useful to construct
SQC .

######  Lemma 3.5.

Let $C$ be a self-dual code of length $n$ over finite chain ring $R$. Then

$\overline{(C:\gamma^{t-1-i})}^{\perp}\subseteq\overline{(C:\gamma^{i+j})},$

where $0\leq i\leq t-1,~{}0\leq j\leq t-1-i$. In particular,

$\overline{(C:\gamma^{t-1})}^{\perp}\subseteq\overline{(C:\gamma^{t-1})}.$

Proof. For $1\leq i\leq t-1,~{}0\leq j\leq t-1-i$, by definition 3.2 and Lemma
3.4, we have

$\overline{(C:\gamma^{t-1-i})}^{\perp}=\overline{(C^{\perp}:\gamma^{i})}=\overline{(C:\gamma^{i})}\subseteq\overline{(C:\gamma^{i+j})}.$

In case $i=0$, obviously,
$\overline{(C:\gamma^{t-1})}^{\perp}\subseteq\overline{(C:\gamma^{t-1})}$. ∎

######  Theorem 3.6.

Let $C$ be a self-dual code of length $n$ and minimum distance $d_{H}(C)$ over
finite chain ring $R$ with a generator matrix $G$ in standard form. Then

$(1)$ there exists a quantum code with parameters $[[n,2k(C)-n,\geq
d_{H}(C)]]_{q}$. In particular, if $C$ is an MDR code, then there exists an
SQCMDS code with parameters $[[n,2k(C)-n,d_{H}(C)]]_{q}$.

$(2)$ there exists a quantum code with parameters $[[n,l+2s-n,\geq
d_{1}]]_{q}$, where
$d_{1}=\mathrm{min}\\{d_{H}\overline{(C:\gamma^{t-1-i})},d_{H}\overline{(C:\gamma^{i+j})}\\}$,
$s=k_{0}(C)+k_{1}(C)+\cdots+k_{i}(C)$,$l=k_{i+1}(C)+\cdots+k_{i+j}(C)$, and
$0\leq i\leq t-1,0\leq j\leq t-1-i$.

Proof. By Lemma 3.4 (1), We know that
$\mathrm{dim}\overline{(C:\gamma^{t-1})}=k(C)$. Thus, there exists a
$[n,k(C),d_{H}(C)]$ code with
$\overline{(C:\gamma^{t-1})}^{\perp}\subseteq\overline{(C:\gamma^{t-1})}$.
According Corollary 2.3, the part (1) is proved.

For (2), by Lemma 3.4 (2),
$\mathrm{dim}\overline{(C:\gamma^{t-1-i})}^{\perp}=k_{0}(C)+k_{1}(C)+\cdots+k_{i}(C)$,
and
$\mathrm{dim}\overline{(C:\gamma^{i+j})}=k_{0}(C)+k_{1}(C)+\cdots+k_{i}(C)+\cdots+k_{i+j}(C)$.
Using Theorem 2.2 and Lemma 3.5, there exists a quantum code with parameters
$[[n,l+2s-n,\geq d_{1}]]_{q}$, which is the required result. ∎

In the rest of this section, we aim to obtain good quantum codes by cyclic
codes over a finite chain ring $\widetilde{R}$ with maximal ideal
$\mathfrak{m}=\widetilde{R}\gamma$, where $\gamma$ is a generator of
$\mathfrak{m}$ with nilpotency index $2$.

The following result is well known (see [11]).

######  Theorem 3.7.

Let $C$ be a cyclic code of length $n$ over finite chain ring $\widetilde{R}$
with characteristic $p^{a}$, where $(p,n)=1$. Then

$(1)$ $C=\langle f(x)h(x),\gamma f(x)g(x)\rangle$, where
$f(x)g(x)h(x)=x^{n}-1$.

$(2)$ $C^{\perp}=\langle g^{*}(x)h^{*}(x),\gamma g^{*}(x)f^{*}(x)\rangle$,
where $g^{*}(x)=x^{\mathrm{deg}g(x)}g(\frac{1}{x})$, i.e., $g^{*}(x)$ is the
reciprocal of $g(x)$.

$(3)$ $\overline{C}=\langle\overline{fh}\rangle$, and
$\overline{(C:\gamma)}=\langle\overline{f}\rangle$.

######  Theorem 3.8.

Let $C$ be a cyclic code of length $n$ over finite chain ring $\widetilde{R}$
with characteristic $p^{a}$, where $(p,n)=1$. If $C=\langle f(x)h(x),\gamma
f(x)g(x)\rangle$ with $f(x)g(x)h(x)=x^{n}-1$, then $C$ is self-dual if and
only if $f(x)=\epsilon g^{*}(x)$ and $h(x)=\varepsilon h^{*}(x)$, where
$\epsilon$ and $\varepsilon$ are units.

Proof. The sufficiency is obvious since $C^{\perp}=\langle
g^{*}(x)h^{*}(x),\gamma g^{*}(x)f^{*}(x)\rangle$.

Now, If $C$ is self-dual, by Theorem 3.7 (2) we know that $\langle
f(x)h(x),\gamma f(x)g(x)\rangle=\langle g^{*}(x)h^{*}(x),\gamma
g^{*}(x)f^{*}(x)\rangle$. But these generators are the unique generators of
this form. Hence

$f(x)h(x)=g^{*}(x)h^{*}(x).$

and

$f(x)h(x)g(x)=g^{*}(x)h^{*}(x)g(x)=-g^{*}(x)h^{*}(x)f^{*}(x)=x^{n}-1.$

Since $f^{*}(x)$ and $g^{*}(x)h^{*}(x)$ are coprime, $f^{*}(x)\mid g(x)$.
Similarly, since

$f(x)h(x)f^{*}(x)=g^{*}(x)h^{*}(x)f^{*}(x)=-f(x)g(x)h(x).$

and $g(x)$ and $f(x)h(x)$ are coprime, $g(x)\mid f^{*}(x)$. That means that
$g(x)=\epsilon f^{*}(x)$. Now, $f(x)h(x)=g^{*}(x)h^{*}(x)=\epsilon
f(x)h^{*}(x)$ where $h^{*}(x)$ and $f(x)$ are coprime. This implies that
$h(x)\mid h^{*}(x)$. Similarly, since $f^{*}(x)h^{*}(x)=g(x)h(x)=\epsilon
f^{*}(x)h(x)$ where $h^{*}(x)$ and $f^{*}(x)$ are coprime, $h^{*}(x)\mid
h(x)$. Therefore, $h(x)=\varepsilon h^{*}(x)$. ∎

Now combining Theorem 3.6 , 3.7 and 3.8, the following result is obtained.

######  Theorem 3.9.

Let $C=\langle f(x)h(x),\gamma f(x)g(x)\rangle$ be a cyclic self-dual code of
length $n$ over finite chain ring $\widetilde{R}$ with characteristic $p^{a}$,
where $(p,n)=1$ and $f(x)g(x)h(x)=x^{n}-1$. Then

$(1)$ there exists a quantum code with parameters
$[[n,n-2\mathrm{deg}\overline{f(x)},\geq d_{H}\overline{(C:\gamma)}~{}]]_{q}$.

$(2)$ there exists a quantum code with parameters
$[[n,n-2\mathrm{deg}\overline{f(x)}-\mathrm{deg}\overline{h(x)},\geq
d_{H}\overline{(C:\gamma)}~{}]]_{q}$.

######  Example 1.

We first provide some examples to obtain quantum codes by non-trivial self-
dual cyclic codes over the chain ring $\mathbb{F}_{2}+u\mathbb{F}_{2}$, where
$u^{2}=0$.

$(1)$Let $n=7$. The factorization of $x^{7}+1$ over
$\mathbb{F}_{2}+u\mathbb{F}_{2}$:

$x^{7}-1=(x+1)(x^{3}+x^{2}+1)(x^{3}+x+1).$

Let $f(x)=x^{3}+x^{2}+1,g(x)=x^{3}+x+1,h(x)=x+1$. Then by Theorem 3.8 cyclic
code $C=\langle f(x)h(x),uf(x)g(x)\rangle$ is self-dual. Using Theorem 3.9
(1), a $[[7,1,\geq 3]]_{2}$ quantum code may be obtained from the
$\overline{(C:u)}=\langle\overline{f}\rangle$ of this code. Obviously, it is a
SQCNMDS code. Again using Theorem 3.9 (2), a $[[7,0,\geq 3]]_{2}$ quantum code
may be obtained from the $\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of
this code.

$(2)$ Let $n=15$. The factorization of $x^{15}+1$ over
$\mathbb{F}_{2}+u\mathbb{F}_{2}$ is equal to$f_{1}f_{2}f_{3}f_{4}f_{5}$, where
$f_{1}=x+1,f_{2}=x^{2}+x+1,f_{3}=x^{4}+x+1,f_{4}=x^{4}+x^{3}+1,f_{5}=x^{4}+x^{3}+x^{2}+x+1$.
Set $f(x)=f_{3},g(x)=f_{4},h(x)=f_{1}f_{2}f_{5}$. Then by Theorem 3.8 cyclic
code $C=\langle f(x)h(x),uf(x)g(x)\rangle$ is self-dual. Using Theorem 3.9
(1), a $[[15,7,\geq 3]]_{2}$ quantum code may be obtained from the
$\overline{(C:u)}=\langle\overline{f}\rangle$ of this code. Again using
Theorem 3.9 (2), a $[[15,0,\geq 3]]_{2}$ quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

$(3)$ Let $n=21$. The factorization of $x^{21}+1$ over
$\mathbb{F}_{2}+u\mathbb{F}_{2}$ is equal to $f_{1}f_{2}f_{3}f_{4}f_{5}f_{6}$,
where
$f_{1}=x+1,f_{2}=x^{2}+x+1,f_{3}=x^{3}+x+1,f_{4}=x^{3}+x^{2}+1,f_{5}=x^{6}+x^{4}+x^{2}+x+1,f_{6}=x^{6}+x^{5}+x^{4}+x^{2}+1$.
Set $f(x)=f_{5},g(x)=f_{6},h(x)=f_{1}f_{2}f_{3}f_{4}$. Then by Theorem 3.8
cyclic code $C=\langle f(x)h(x),uf(x)g(x)\rangle$ is self-dual. Using Theorem
3.9 (1), a $[[21,9,\geq 3]]_{2}$ quantum code may be obtained from the
$\overline{(C:u)}=\langle\overline{f}\rangle$ of this code. Again using
Theorem 3.9 (2), a $[[21,0,\geq 3]]_{2}$ quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

$(4)$ Let $n=23$. The factorization of $x^{23}+1$ over
$\mathbb{F}_{2}+u\mathbb{F}_{2}$ is equal to $f_{1}f_{2}f_{3}$, where
$f_{1}=x+1,f_{2}=x^{11}+x^{9}+x^{7}+x^{6}+x^{5}+x+1,f_{3}=x^{11}+x^{10}+x^{6}+x^{5}+x^{4}+x^{2}+1$.
Set $f(x)=f_{2},g(x)=f_{3},h(x)=f_{1}$. Then by Theorem 3.8 cyclic code
$C=\langle f(x)h(x),uf(x)g(x)\rangle$ is self-dual. Using Theorem 3.9 (1), a
$[[23,1,\geq 7]]_{2}$ quantum code may be obtained from the
$\overline{(C:u)}=\langle\overline{f}\rangle$ of this code. Again using
Theorem 3.9 (2), a $[[23,0,\geq 7]]_{2}$ quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

$(5)$ Let $n=31$. The factorization of $x^{31}+1$ over
$\mathbb{F}_{2}+u\mathbb{F}_{2}$ is equal to
$(x+1)f_{1}f_{1c}f_{2}f_{2c}f_{3}f_{3c}$, where
$f_{1}=x^{5}+x^{2}+1,f_{1c}=x^{5}+x^{3}+1,f_{2}=x^{5}+x^{3}+x^{2}+x+1,f_{2c}=x^{5}+x^{4}+x^{3}+x^{2}+1,f_{3}=x^{5}+x^{4}+x^{2}+x+1,f_{3c}=x^{5}+x^{4}+x^{3}+x+1$.
Set $f(x)=f_{1}f_{2},g(x)=f_{1c}f_{2c},h(x)=(x+1)f_{3}f_{3c}$. Then by Theorem
3.8 cyclic code $C=\langle f(x)h(x),uf(x)g(x)\rangle$ is self-dual. Using
Theorem 3.9 (1), a $[[31,21,\geq 5]]_{2}$ quantum code may be obtained from
the $\overline{(C:u)}=\langle\overline{f}\rangle$ of this code. Obviously, it
is a SQCNMDS code. Again using Theorem 3.9 (2), a $[[31,0,\geq 5]]_{2}$
quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

######  Example 2.

We provide some examples to obtain quantum codes by non-trivial cyclic self-
dual codes over the chain ring $\mathbb{Z}_{p^{2}}$.

$(1)$ Length $11$ over $\mathbb{Z}_{3^{2}}$. The factorization of $x^{11}-1$
over $\mathbb{Z}_{3^{2}}$ into a product of basic irreducible polynomials over
$\mathbb{Z}_{3^{2}}$ is given by

$x^{11}-1=f_{1}(x)f_{2}(x)f_{3}(x).$

where
$f_{1}=x-1,f_{2}=x^{5}+3x^{4}+8x^{3}+x^{2}+2x-1,f_{3}=x^{5}-2x^{4}-x^{3}+x^{2}-3x-1$.
Let $f(x)=-f_{3}(x),g(x)=f_{1}(x),h(x)=-f_{1}(x)$. Then by Theorem 3.8 cyclic
code $C=\langle f(x)h(x),3f(x)g(x)\rangle$ is self-dual. Using Theorem 3.9
(1), a $[[11,1,\geq 2]]_{3}$ quantum code may be obtained from the
$\overline{(C:3)}=\langle\overline{f}\rangle$ of this code. Again using
Theorem 3.9 (2), a $[[11,0,\geq 2]]_{3}$ quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

Length $13$ over $\mathbb{Z}_{3^{2}}$. The factorization of $x^{13}-1$ over
$\mathbb{Z}_{3^{2}}$ into a product of basic irreducible polynomials over
$\mathbb{Z}_{3^{2}}$ is given by

$x^{13}-1=f_{1}(x)f_{2}(x)f_{3}(x)f_{4}(x)f_{5}(x).$

where
$f_{1}=x-1,f_{2}=x^{3}+6x^{2}+2x+8,f_{3}=x^{3}+7x^{2}+3x+8,f_{4}=x^{3}+4x^{2}+7x+8,f_{5}=x^{3}+2x^{2}+7x+8$.
Let $f(x)=f_{3}(x),g(x)=f_{2}(x),h(x)=-f_{1}(x)f_{4}(x)f_{5}(x)$. Then by
Theorem 3.8 cyclic code $C=\langle f(x)h(x),3f(x)g(x)\rangle$ is self-dual.
Using Theorem 3.9 (1), a $[[13,7,\geq 3]]_{3}$ quantum code may be obtained
from the $\overline{(C:3)}=\langle\overline{f}\rangle$ of this code.
Obviously, it is a SQCNMDS code. Again using Theorem 3.9 (2), a $[[13,0,\geq
3]]_{3}$ quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

$(2)$ Length $11$ over $\mathbb{Z}_{5^{2}}$. The factorization of $x^{11}-1$
over $\mathbb{Z}_{5^{2}}$ into a product of basic irreducible polynomials over
$\mathbb{Z}_{5^{2}}$ is given by

$x^{11}-1=f_{1}(x)f_{2}(x)f_{3}(x).$

where
$f_{1}=x-1,f_{2}=x^{5}+17x^{4}+24x^{3}+x^{2}+16x+24,f_{3}=x^{5}+9x^{4}+24x^{3}+x^{2}+8x+24$.
Let $f(x)=-f_{3}(x),g(x)=f_{2}(x),h(x)=-f_{1}(x)$. Then by Theorem 3.8 cyclic
code $C=\langle f(x)h(x),5f(x)g(x)\rangle$ is self-dual. Using Theorem 3.9
(1), a $[[11,1,\geq 5]]_{5}$ quantum code may be obtained from the
$\overline{(C:5)}=\langle\overline{f}\rangle$ of this code. Obviously, it is a
SQCNMDS code. Again using Theorem 3.9 (2), a $[[11,0,\geq 5]]_{5}$ quantum
code may be obtained from the $\overline{C}=\langle\overline{f(x)h(x)}\rangle$
of this code.

Length $22$ over $\mathbb{Z}_{5^{2}}$. The factorization of $x^{22}-1$ over
$\mathbb{Z}_{5^{2}}$ into a product of basic irreducible polynomials over
$\mathbb{Z}_{5^{2}}$ is given by

$x^{22}-1=f_{1}(x)f_{2}(x)f_{3}(x)f_{4}(x)f_{5}(x)f_{6}(x).$

where
$f_{1}=x-1,f_{2}=x+24,f_{3}=x^{5}+16x^{4}+24x^{3}+24x^{2}+8x+1,f_{4}=x^{5}+17x^{4}+24x^{3}+x^{2}+16x+1,f_{5}=x^{5}+8x^{4}+24x^{3}+24x^{2}+16x+1,f_{6}=x^{5}+9x^{4}+24x^{3}+x^{2}+16x+1$.
Let $f(x)=f_{5}(x),g(x)=f_{3}(x),h(x)=f_{1}(x)f_{2}(x)f_{4}(x)f_{6}(x)$. Then
by Theorem 3.8 cyclic code $C=\langle f(x)h(x),5f(x)g(x)\rangle$ is self-dual.
Using Theorem 3.9 (1), a $[[22,12,\geq 2]]_{5}$ quantum code may be obtained
from the $\overline{(C:5)}=\langle\overline{f}\rangle$ of this code. Again
using Theorem 3.9 (2), a $[[22,0,\geq 3]]_{5}$ quantum code may be obtained
from the $\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

$(3)$ Length $6$ over $\mathbb{Z}_{7^{2}}$. The factorization of $x^{6}-1$
over $\mathbb{Z}_{7^{2}}$ into a product of basic irreducible polynomials over
$\mathbb{Z}_{7^{2}}$ is given by

$x^{6}-1=f_{1}(x)f_{2}(x)f_{3}(x)f_{4}(x)f_{5}(x)f_{6}(x).$

where
$f_{1}=x-1,f_{2}=x+1,f_{3}=x+18,f_{4}(x)=x-18,f_{5}(x)=x+19,f_{6}(x)=x-19$.
Let $f(x)=-18f_{6}(x),g(x)=f_{4}(x),h(x)=-18f_{1}(x)f_{2}(x)f_{3}(x)f_{5}(x)$.
Then by Theorem 3.8 cyclic code $C=\langle f(x)h(x),7f(x)g(x)\rangle$ is self-
dual. Using Theorem 3.9 (1), a $[[6,4,\geq 2]]_{7}$ quantum code may be
obtained from the $\overline{(C:7)}=\langle\overline{f}\rangle$ of this code.
Obviously, it is a SQCMDS code. Again using Theorem 3.9 (2), a $[[6,0,\geq
2]]_{7}$ quantum code may be obtained from the
$\overline{C}=\langle\overline{f(x)h(x)}\rangle$ of this code.

##  4 SQC from to the Cyclic Codes over Chain Rings
$\mathbb{F}_{p^{2m}}+u\mathbb{F}_{p^{2m}}$

Throughout this section, $p$ denotes a prime number and $\mathbb{F}_{p^{2m}}$
denotes the finite field with $p^{2m}$ elements for a positive integer $m$. We
always assume that $n$ is a positive integer.

The ring $R=\mathbb{F}_{p^{2m}}+u\mathbb{F}_{p^{2m}}$ consists of all
$p^{2m}$-ary polynomials of degree $0$ and $1$ in an indeterminate $u$, and it
is closed under $p^{2m}$-ary polynomial addition and multiplication modulo
$u^{2}$. Thus, $R=\frac{\mathbb{F}_{p^{2m}}[u]}{\langle
u^{2}\rangle}=\\{a+ub|a,b\in\mathbb{F}_{p^{2m}}\\}$ is a local ring with
maximal ideal $u\mathbb{F}_{p^{2m}}$. Therefore, it is a chain ring. The ring
$R$ has precisely $p^{2m}(p^{2m}-1)$ units, which are of the forms
$\alpha+u\beta$ and $\gamma$, where $\alpha,\beta$, and $\gamma$ are nonzero
elements of the field $\mathbb{F}_{p^{2m}}$.

Let $\overline{a+ub}:=\overline{a}+u\overline{b}$, where
$\overline{a}=a^{p^{m}}$, and $\overline{b}=b^{p^{m}}$. The Hermitian inner
product over $\mathbb{F}_{p^{2m}}+u\mathbb{F}_{p^{2m}}$ is defined as follows:

$[\mathbf{x},\mathbf{y}]_{H}=\sum_{i=1}^{n}x_{i}\overline{y_{i}}.$

where $\mathbf{x},\mathbf{y}\in R^{n}$, $\mathbf{x}=(x_{1},\ldots,x_{n})$ and
$\mathbf{y}=(y_{1},\ldots,y_{n})$. The Hermitian dual code $C^{\perp_{H}}$ of
$C$ is defined by

$C^{\perp_{H}}=\\{\mathbf{x}\in
R^{n}~{}|~{}[\mathbf{x},\mathbf{y}]_{H}=0~{}\mathrm{for~{}all}~{}~{}\mathbf{y}~{}\in
C\\}.$

It is evident that $C^{\perp_{H}}$ is linear. We say that a code $C$ is
Hermitian dual-containing code if $C^{\perp_{H}}\subset C$ and $C\neq R^{n}$,
and Hermitian self-dual if $C^{\perp_{H}}=C$.

It is easy to prove that

$~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}\mid C\mid\cdot\mid
C^{\perp_{H}}\mid=\mid
R\mid^{n}.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.1)$

The following lemma can be found in [12].

######  Lemma 4.1.

$[12,~{}Corollary~{}4.2]$ Assume the notations given above. Then there exist
$\alpha\in R$ such that $\alpha^{2}=-1$ if and only if $p^{2m}\equiv
1~{}(\mathrm{mod}~{}4)$.

######  Remark 4.2.

Since $\alpha\in R$, there exist $s,t\in\mathbb{F}_{p^{2m}}$ such that
$\alpha=s+ut$. Hence computing in $R$, we have $\alpha^{2}=s^{2}+2stu=-1$,
which implies that $s^{2}=-1,2st=0$. If $p=2$, then take
$s=1\in\mathbb{F}_{p^{2m}},t=0$ we have $\alpha^{2}=-1$; If $p\neq 2$, then
$t=0$ since $2st=0$. Therefore, $\alpha=s\in\mathbb{F}_{p^{2m}}$.

From now on, we always assume that $p^{m}\equiv 1~{}(\mathrm{mod}~{}4)$, then
$p^{2m}\equiv 1~{}(\mathrm{mod}~{}4)$. So there exist
$\alpha\in\mathbb{F}_{p^{2m}}$ such that $\alpha^{2}=-1$ in $R$.

We first give the definition of the Gray map on $R^{n}$. The Gray map
$\Phi_{1}:~{}R\rightarrow\mathbb{F}_{p^{2m}}^{2}$ is given by
$\Phi_{1}(a+bu)=(\alpha b,a+b)$, where $\alpha^{2}=-1$. This map can be
extended to $R^{n}$ in a natural way:

$\displaystyle\Phi:$ $\displaystyle
R^{n}\longrightarrow\mathbb{F}_{p^{2m}}^{2n}$
$\displaystyle(a_{1}+ub_{1},\ldots,a_{n}+ub_{n})\longmapsto(\alpha
b_{1},a_{1}+b_{1},\ldots,\alpha b_{n},a_{n}+b_{n}).$

Next, we define a Gray weight for codes over $R$ as follows.

######  Definition 4.3.

The Gray weight over $R$ is a weight function on $R$ defined as:

$w_{G}(a+bu)=\left\\{\begin{aligned}
&0~{}~{}~{}~{}~{}~{}~{}~{}~{}if~{}a=0,~{}b=0,\\\
&1~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}if~{}a\neq 0,~{}b=0,\\\
&1~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}if~{}b\neq 0,a+b\equiv
0~{}(\mathrm{mod~{}}p),\\\ &2~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}if~{}b\neq
0,a+b\neq 0~{}(\mathrm{mod~{}}p).\end{aligned}\right.$

Define the Gray weight of a codeword $\mathbf{c}=(c_{1},\ldots,c_{n})\in
R^{n}$ to be the rational sum of the Gray weight of its components is,
$w_{G}(\mathbf{c})=\sum_{i=1}^{n}w_{G}(c_{i})$. For any
$\mathbf{c}_{1},\mathbf{c}_{2}\in R^{n}$, the Gray distance $d_{G}$ is given
by
$d_{G}(\mathbf{c}_{1},\mathbf{c}_{2})=w_{G}(\mathbf{c}_{1}-\mathbf{c}_{2})$.
The minimum Gray distance of $C$ is the smallest nonzero Gray distance between
all pairs of distinct codewords of $C$. The minimum Gray weight of $C$ is the
smallest nonzero Gray weight among all codewords of $C$. If $C$ is linear,
then the minimum Gray distance is same as the minimum Gray weight.

The Hamming weight $W(\mathbf{c})$ of a codeword $\mathbf{c}$ is the number of
nonzero components in $\mathbf{c}$. The Hamming distance
$d(\mathbf{c}_{1},\mathbf{c}_{2})$ between two codewords $\mathbf{c}_{1}$ and
$\mathbf{c}_{2}$ is Hamming weight of the codeword
$\mathbf{c}_{1}-\mathbf{c}_{2}$. The minimum Hamming distance $d$ of $C$ is
defined as
$\mathrm{min}\\{d(\mathbf{c}_{1},\mathbf{c}_{2})|\mathbf{c}_{1},\mathbf{c}_{2}\in
C,\mathbf{c}_{1}\neq\mathbf{c}_{2}\\}$ (See $\mathrm{[13]}$).

######  Proposition 4.4.

The Gray map $\Phi$ is a distance-preserving map from
$(R^{n},~{}\mathrm{Gray~{}distance})$ to
$(\mathbb{F}_{p^{2m}}^{2n},~{}\mathrm{Hamming~{}distance})$ and it is also
$\mathbb{F}_{p^{2m}}$-linear.

Proof. From the above definitions, it is clear that
$\Phi(\mathbf{c}_{1}-\mathbf{c}_{2})=\Phi(\mathbf{c}_{1})-\Phi(\mathbf{c}_{2})$
for $\mathbf{c}_{1},\mathbf{c}_{2}\in R^{n}$. Thus,

$d_{G}(\mathbf{c}_{1},\mathbf{c}_{2})=w_{G}(\mathbf{c}_{1}-\mathbf{c}_{2})=w(\Phi(\mathbf{c}_{1}-\mathbf{c}_{2}))=w(\Phi(\mathbf{c}_{1})-\Phi(\mathbf{c}_{2}))=d(\Phi(\mathbf{c}_{1}),\Phi(\mathbf{c}_{2})).$

Let $\mathbf{c}_{1},\mathbf{c}_{2}\in R^{n}$,
$k_{1},k_{2}\in\mathbb{F}_{p^{2m}}$. From the definition of the Gray map, we
have
$\Phi(k_{1}\mathbf{c}_{1}+k_{2}\mathbf{c}_{2})=k_{1}\Phi(\mathbf{c}_{1})+k_{2}\Phi(\mathbf{c}_{2})$,
which means that $\Phi$ is an $\mathbb{F}_{p^{2m}}$-linear map. ∎

######  Corollary 4.5.

If $C$ is a linear code over $R$ of length $n$, size $(p^{2m})^{k}$ and
minimum Gray weight $d_{G}$, then $\Phi(C)$ is a linear code over
$\mathbb{F}_{p^{2m}}$ with parameters $[2n,k,d_{G}]$.

The Hermitian inner product over $\mathbb{F}_{p^{2m}}$ is defined as follows:

$[\mathbf{a},\mathbf{b}]_{H}=\mathbf{a}\cdot\overline{\mathbf{b}}=\sum_{i=1}^{n}a_{i}b_{i}^{p^{m}},$

where
$\mathbf{a},\mathbf{b}\in\mathbb{F}_{p^{2m}}^{n},\mathbf{a}=(a_{1},\ldots,a_{n}),\mathbf{b}=(b_{1},\ldots,b_{n})$
and $\cdot$ is the usual Euclidean inner product.

An important connection that we want to investigate is the relation between
the Hermitian dual and the Gray image of a code. We have the following
theorem.

######  Theorem 4.6.

Let $C$ be a linear code over $R$ of length $n$. Then
$\Phi(C^{\perp_{H}})=\Phi(C)^{\perp_{H}}$.

Proof. To prove the theorem, we first show
$\Phi(C^{\perp_{H}})\subset\Phi(C)^{\perp_{H}}$, i.e.,

$[\mathbf{x},\mathbf{y}]_{H}=0~{}\Rightarrow~{}[\Phi(\mathbf{x}),\Phi(\mathbf{y})]_{H}=0~{}~{}\mathrm{for~{}all}~{}\mathbf{x},\mathbf{y}\in
R^{n}.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.2)$

To this extent, let us assume that
$\mathbf{x}=(a_{1}+ub_{1},\ldots,a_{n}+ub_{n})$ and
$\mathbf{y}=(c_{1}+ud_{1},\ldots,c_{n}+ud_{n})$, where
$a_{i},b_{i},c_{i},d_{i}\in\mathbb{F}_{p^{2m}}$. Then by

$[\mathbf{x},\mathbf{y}]_{H}=\sum_{i=1}^{n}a_{i}\overline{c_{i}}+\sum_{i=1}^{n}(b_{i}\overline{c_{i}}+a_{i}\overline{d_{i}})u,$

we see that $[\mathbf{x},\mathbf{y}]_{H}=0$ if and only if

$~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}\sum_{i=1}^{n}a_{i}\overline{c_{i}}=0,~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.3)$

and

$~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}\sum_{i=1}^{n}(b_{i}\overline{c_{i}}+a_{i}\overline{d_{i}})=0.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.4)$

Note that $p^{m}\equiv 1~{}(\mathrm{mod}~{}4)$ we can assume that $p^{m}=4k+1$
for some $k\in\mathbb{N}$, hence, $p^{m}+1=4k+2=2(2k+1)$. According to
$\alpha^{2}=-1$, we have

$~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}\alpha^{p^{m}+1}=(\alpha^{2})^{2k+1}=-1.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.5)$

Now, since $\Phi(\mathbf{x})=(\alpha b_{1},a_{1}+b_{1},\ldots,\alpha
b_{n},a_{n}+b_{n})$, and $\Phi(\mathbf{y})=(\alpha
d_{1},c_{1}+d_{1},\ldots,\alpha d_{n},c_{n}+d_{n})$, we get

$[\Phi(\mathbf{x}),\Phi(\mathbf{y})]_{H}=\sum_{i=1}^{n}\alpha^{p^{m}+1}b_{i}\overline{d_{i}}+\sum_{i=1}^{n}(a_{i}+b_{i})(\overline{c_{i}+d_{i}})~{}~{}~{}~{}~{}~{}$
$~{}~{}~{}~{}~{}=\sum_{i=1}^{n}\alpha^{p^{m}+1}b_{i}\overline{d_{i}}+\sum_{i=1}^{n}(a_{i}\overline{c_{i}}+a_{i}\overline{d_{i}}+b_{i}\overline{c_{i}}+b_{i}\overline{d_{i}})$
$~{}~{}~{}~{}~{}~{}~{}~{}=\sum_{i=1}^{n}(\alpha^{p^{m}+1}+1)b_{i}\overline{d_{i}}+\sum_{i=1}^{n}a_{i}\overline{c_{i}}+\sum_{i=1}^{n}(b_{i}\overline{c_{i}}+a_{i}\overline{d_{i}}),$

by(4.3-4.5) which finishes the of (4.2), i.e.,

$\Phi(C^{\perp_{H}})\subset\Phi(C)^{\perp_{H}}.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.6)$

In the light of Corollary 4.5, $\Phi(C)$ is a linear code of length $2n$ of
size $|C|$ over $\mathbb{F}_{p^{2m}}$. So, by Corollary 4.5, we know that

$|\Phi(C)^{\perp_{H}}|=\frac{(p^{2m})^{2n}}{|\Phi(C)|}=\frac{(p^{2m})^{2n}}{|C|}.$

Since $R$ is a finite chain ring, i.e., Frobenius ring, we have

$|C^{\perp_{H}}|\cdot|C|=|R|^{n}=(p^{2m})^{2n}.$

Hence, this implies that

$|\Phi(C^{\perp_{H}})|=|\Phi(C)^{\perp_{H}}|.~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}(4.6)$

Combining (4.5) with (4.6), we get the desired equality. ∎

The following is an immediate corollary to this:

######  Corollary 4.7.

(1) If $C$ is a Hermitian self-dual code of length $n$ over $R$, then
$\Phi(C)$ is a Hermitian self-dual code of length $2n$ over
$\mathbb{F}_{p^{2m}}$;

(2) If $C$ is a Hermitian dual-containing code of length $n$ over $R$, then
$\Phi(C)$ is a Hermitian dual-containing code of length $2n$ over
$\mathbb{F}_{p^{2m}}$.

In the folowing, we always assume that $n$ is a positive integer and
$(n,p)=1$. Let $\mathcal{R}_{n}=\frac{R[x]}{<x^{n}-1>}$. We denote by $\mu$
the natural surjective ring morphism from $R$ to $\mathbb{F}_{p^{2m}}$, which
can be extended naturally to a surjective ring morphism from $R[x]$ to
$\mathbb{F}_{p^{2m}}[x]$.

For a polynomial $f(x)$ of degree $k$ in $R[x]$, its reciprocal polynomial
$x^{k}f(x^{-1})$ is denoted by $f^{*}(x)$. Note that the roots of $f^{*}(x)$
are the reciprocal of the corresponding roots of $f(x)$. Set
$f(x)=a_{0}+a_{1}x+\cdots+a_{k}x^{k}$, we define

$\overline{f(x)}=\overline{a}_{0}+\overline{a}_{1}x+\cdots+\overline{a}_{k}x^{k}.$

######  Lemma 4.8.

Let $f(x)=(t_{0}+us_{0})+(t_{1}+us_{1})x+\cdots+(t_{n-1}+us_{n-1})x^{n-1}\in
R[x]$ and $\eta$ be a primitive $n$th root of unity in some extension ring of
$R$. If $\eta^{s}$ is a root of $f(x)$, there $\overline{f^{*}(x)}$ has
$\eta^{-p^{m}s}$ as a root.

Proof. Let $\xi=\eta^{s}$. Then

$f(\xi)=(t_{0}+t_{1}\xi+\cdots+t_{n-1}\xi^{n-1})+u(s_{0}+s_{1}\xi+\cdots+s_{n-1}\xi^{n-1})=0.$

This implies

$t_{0}+t_{1}\xi+\cdots+t_{n-1}\xi^{n-1}=0~{}~{}\mathrm{and}~{}~{}s_{0}+s_{1}\xi+\cdots+s_{n-1}\xi^{n-1}=0.$

Therefore,

$\overline{f^{*}(\xi^{-p^{m}})}=(\overline{t}_{n-1}+\overline{t}_{n-2}\xi^{-p^{m}}+\cdots+\overline{t}_{0}\xi^{-p^{m}(n-1)})+u(\overline{s}_{n-1}+\overline{s}_{n-2}\xi^{-p^{m}}+\cdots+\overline{s}_{0}\xi^{-p^{m}(n-1)})$
$=\xi^{-p^{m}(n-1)}[(t_{0}+t_{1}\xi+\cdots+t_{n-1}\xi^{n-1})^{p^{m}}+u(s_{0}+s_{1}\xi+\cdots+s_{n-1}\xi^{n-1})^{p^{m}}]=0$

∎

Let $i$ be an integer such that $0\leq i\leq n-1$, and let $l$ be the smallest
positive integer such that $i(p^{2m})^{l}\equiv i~{}\mathrm{(mod}~{}n)$. Then
$C_{i}=\\{i,ip^{2m},\ldots,i(p^{2m})^{l-1}\\}$ is the $p^{2m}-$cyclotomic
coset module $n$ containing $i$. A cyclotomic coset $C_{i}$ is called
symmetric if $n-p^{m}i\in C_{i}$ and asymmetric otherwise. Let $I_{1}$ and
$I_{2}$ be sets of symmetric and asymmetric coset representatives modulo $n$,
respectively. Since $p$ is coprime with $n$, the irreducible factors of
$x^{n}-1$ in $\mathbb{F}_{p^{2m}}[x]$ can be described by the
$p^{2m}-$cyclotomic cosets. Suppose that $\zeta$ be a primitive $n$th root of
unity over some extension field of $\mathbb{F}_{p^{2m}}$. Then $\zeta$ is also
a primitive $n$th root of unity over some extension ring of $R$. Let
$m_{j}(x)$ be the minimal polynomial of $\zeta^{j}$ with respect to
$\mathbb{F}_{p^{2m}}$. Then $m_{j}(x)=\Pi_{i\in C_{j}}(x-\zeta^{i})$, and
$\overline{m_{j}^{*}(x)}=\Pi_{i\in C_{-p^{m}j}}(x-\zeta^{i})$ by Lemma 4.8.
Therefore, polynomial $x^{n}-1$ factors uniquely into monic irreducible
polynomial in $\mathbb{F}_{p^{2m}}[x]$ as $x^{n}-1=\Pi_{j\in
I_{1}}m_{j}(x)\Pi_{j\in I_{2}}m_{j}(x)m_{-p^{m}j}(x)$. By
$\mathrm{Hensel}^{,}$s lemma (See [9,Theorem 4.1.1]), $x^{n}-(1+u)$ has a
unique decomposition as a product $\Pi_{j\in I_{1}}M_{j}(x)\Pi_{j\in
I_{2}}M_{j}(x)M_{-p^{m}j}(x)$ of pairwise coprime monic basic irreducible
polynomials in $R[x]$ with $\mu(M_{j}(x))=m_{j}(x)$ for each $j\in I_{1}\cup
I_{2}$.

The following two lemmas can be found in [14].

######  Lemma 4.9.

$[14,~{}Theorem~{}3.4]$ Let $x^{n}-(1+u)=\Pi_{j\in I_{1}\cup I_{2}}M_{j}(x)$
be the unique factorization of $x^{n}-(1+u)$ into a product of monic basic
irreducible pairwise coprime polynomials in $R[x]$. If $C$ is a cyclic code of
length $n$ over $R$, then $C=\langle\Pi_{j\in I_{1}\cup
I_{2}}M_{j}^{k_{j}}(x)\rangle$, where $0\leq k_{j}\leq 2$. In this case, $\mid
C\mid=(p^{2m})^{\sum_{j\in I_{1}\cup I_{2}}(2-k_{j})~{}deg~{}M_{j}}$.

######  Lemma 4.10.

$[14,~{}Lemma~{}4.2]$ Let $C=\langle\Pi_{j\in I_{1}\cup
I_{2}}M_{j}^{k_{j}}(x)\rangle$ be a cyclic code of length $n$ over $R$, where
the polynomials $M_{j}$ are the pairwise coprime monic basic irreducible
factors of $x^{n}-(1+u)$ in $R[x]$ and $0\leq k_{j}\leq 2$ for each $j\in I$.
Then $C^{\perp_{H}}=\langle\Pi_{j\in I_{1}\cup
I_{2}}\overline{M^{*}_{j}(x)}^{2-k_{j}}(x)\rangle$ and $\mid
C^{\perp_{H}}\mid=(p^{2m})^{\sum_{j\in I_{1}\cup I_{2}}k_{j}~{}deg~{}M_{j}}$.

######  Theorem 4.11.

Let $C$ be a cyclic code of length $n$ over $R$. If

$C=<\Pi_{j\in I_{1}}M_{j}^{k_{j}}(x)\Pi_{j\in
I_{2}}M_{j}^{i_{j}}(x)M_{-p^{m}j}^{l_{j}}(x)>,$

then $C^{\perp_{H}}\subset C$ if and only if $k_{j}=0$ or $k_{j}=1$ for $j\in
I_{1}$ and $i_{j}+l_{j}\leq 2$ for $j\in I_{2}$.

Proof. According to Lemma 4.10, we have

$C^{\perp_{H}}=\langle\Pi_{j\in I_{1}}M_{j}^{2-k_{j}}(x)\Pi_{j\in
I_{2}}M_{j}^{2-l_{j}}(x)M_{-p^{m}j}^{2-i_{j}}(x)\rangle.$

Comparing with $C=\langle\Pi_{j\in I_{1}}M_{j}^{k_{j}}(x)\Pi_{j\in
I_{2}}M_{j}^{i_{j}}(x)M_{-p^{m}j}^{l_{j}}(x)\rangle$, it follows that
$C^{\perp_{H}}\subset C$ if and only if $k_{j}=0$ or $k_{j}=1$ for $j\in
I_{1}$ and $i_{j}+l_{j}\leq 2$ for $j\in I_{2}$. ∎

From now on, we always assume that $n=sp^{t}-1$. Obviously, $(n,p)=1$. In this
case, we give a method to decompose $x^{n}-(1+u)$ into monic basic irreducible
polynomials in $R(x)$. Let $g_{1}(x),g_{2}(x),\ldots,g_{r}(x)$ be monic basic
irreducible polynomials in $R(x)$ such that $x^{n}-1=g_{1}(x)g_{2}(x)\cdots
g_{r}(x)$. Note that $(1+u)^{p}=1$ and $(1+u)^{sp^{t}}=1$, we have

$(1+u)^{n}(x^{n}-(1+u))=(1+u)^{\sum_{i=1}^{r}deg~{}g_{i}}[(1+u)^{p-deg~{}g_{1}}g_{1}((1+u)x)]\cdots[(1+u)^{p-deg~{}g_{r}}g_{r}((1+u)x)]$
$~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}~{}=(1+u)^{n}[(1+u)^{p-deg~{}g_{1}}g_{1}((1+u)x)]\cdots[(1+u)^{p-deg~{}g_{r}}g_{r}((1+u)x)].$

Therefore,

$x^{n}-(1+u)=[(1+u)^{p-deg~{}g_{1}}g_{1}((1+u)x)]\cdots[(1+u)^{p-deg~{}g_{r}}g_{r}((1+u)x)].$

Let $f_{i}(x)=(1+u)^{p-deg~{}g_{i}}g_{i}((1+u)x)$ for $1\leq i\leq r$. Then
the polynomial $x^{n}-(1+u)$ factor uniquely into monic basic irreducible
polynomials in $R(x)$ as $f_{1}(x)f_{2}(x)\cdots f_{r}(x)$.

For a code $C$ of length $n$ over $R$, their torsion and residue codes are
codes over $\mathbb{F}_{p^{2m}}$, defined as follows.

$\mathrm{Tor}(C)=\\{a\in\mathbb{F}_{p^{2m}}^{n}\mid ua\in
C\\},~{}~{}~{}~{}\mathrm{Res}(C)=\\{a\in\mathbb{F}_{p^{2m}}^{n}\mid\exists
b\in\mathbb{F}_{p^{2m}}^{n}:a+ub\in C\\}.$

The reduction modulo $u$ from $C$ to $\mathrm{Res}(C)$ is given by
$\varphi:C\rightarrow\mathrm{Res}(C),\varphi(a+ub)=a$. Clearly, $\varphi$ is
well defined and onto, with $\mathrm{Ker}(\varphi)=u\mathrm{Tor}(C)$, and
$\varphi(C)=\mathrm{Res}(C)$. Therefore,
$|\mathrm{Res}(C)|=\frac{|C|}{|\mathrm{Tor}(C)|}$. Thus, we obtain
$|C|=|\mathrm{Res}(C)||\mathrm{Tor}(C)|$ and the following two theorems.

######  Theorem 4.12.

Let $C=\langle\Pi_{j\in I}M_{j}^{k_{j}}(x)\rangle$ be a cyclic code of length
$n$ over $R$ where $x^{n}-(1+u)=\Pi_{j\in I_{1}}M_{j}(x)\Pi_{j\in
I_{2}}M_{j}(x)M_{-p^{m}j}(x)$, $0\leq k_{j}\leq 2$, and $I=I_{1}\cup I_{2}$.
Then

$(1)$ $\mathrm{Res}(C)=\langle\Pi_{j\in
I}[\varphi(M_{j}(x))]^{\delta_{j}}\rangle$, where $\delta_{j}=k_{j}$ if
$k_{j}=1$ or $0$, and $\delta_{j}=1$ if $k_{j}=2$;

$(2)$ $\mathrm{Tor}(C)=\langle\Pi_{j\in I}[(\varphi
M_{j}(x))]^{\eta_{j}}\rangle$, where $\eta_{j}=0$ if $k_{j}=1$ or $0$, and
$\eta_{j}=1$ if $k_{j}=2$.

Proof. According to the definition of $\mathrm{Res}(C)$, we have
$\mathrm{Res}(C)=\langle\Pi_{j\in I}[(\varphi M_{j}(x))]^{k_{j}}\rangle$. Note
that if $f(x)$ is a monic irreducible divisor of $x^{n}-1$ in
$\mathbb{F}_{p^{2m}}$ and $g(x)=\frac{x^{n}-1}{f(x)}$, then $(f(x),g(x))=1$.
So there exist $a(x),b(x)\in\mathbb{F}_{p^{2m}}[x]$ such that
$a(x)f(x)+b(x)g(x)=1$ in $\mathbb{F}_{p^{2m}}[x]$. Computing in
$\frac{\mathbb{F}_{p^{2m}}[x]}{\langle x^{n}-1\rangle}$, we get

$a(x)f^{2}(x)=(1-b(x)g(x))f(x)=f(x)-b(x)f(x)g(x)=f(x)-b(x)(x^{n}-1)=f(x).$

Consequently, $\langle f^{2}(x)\rangle=\langle f(x)\rangle$. This proves the
(1).

For (2), since $u\Pi_{j\in I}[(\varphi M_{j}(x))]^{\eta_{j}}=u\Pi_{j\in
I}[M_{j}(x)]^{\eta_{j}}=-\Pi_{j\in I}[M_{j}(x)]^{\eta_{j}+1}\in C$, we have
$\langle\Pi_{j\in I}[(\varphi
M_{j}(x))]^{\eta_{j}}\rangle\subset\mathrm{Tor}(C)$. By Lemma 4.9.3 and
$|C|=|\mathrm{Res}(C)||\mathrm{Tor}(C)|$, we imply

$|\langle\Pi_{j\in I}[(\varphi
M_{j}(x))]^{\eta_{j}}\rangle\mid=\mid\mathrm{Tor}(C)|.$

Thus, $\mathrm{Tor}(C)=\langle\Pi_{j\in I}[(\varphi
M_{j}(x))]^{\eta_{j}}\rangle$. ∎

######  Theorem 4.13.

Let $C$ be a cyclic code of length $n$ over $R$, and let $d_{1}$ and $d_{2}$
be the minimum Hamming distances of the $\mathrm{Res}(C)$ and
$\mathrm{Tor}(C)$, respectively. Then
$d_{G}(C)=\mathrm{min}\\{d_{1},2d_{2}\\}$.

Proof. For any nonzero codeword $c=a(x)+ub(x)\in C$, if $a(x)\neq 0$, then
$a(x)\in\mathrm{Res}(C)$. Thus $w_{G}(c)\geq d_{1}$. Otherwise,
$c=ub(x)\in\mathrm{Tor}(C)$, hence, $w_{G}(c)\geq 2d_{2}$. So
$d_{G}(C)\geq\mathrm{min}\\{d_{1},2d_{2}\\}$. On the other hand, since
$u\mathrm{Tor}(C)$ is contained in $C$, we can obtain $d_{G}(C)\leq 2d_{2}$.
Obviously, $\mathrm{Res}(C)\subset C$, hence $d_{1}\geq d_{G}(C)$. It follows
that $\mathrm{min}\\{d_{1},2d_{2}\\}\geq d_{G}(C)$. This proves the expected
result. ∎

Combing Corollary 2.3, 4.5, 4.7 and Theorem 4.13, we have the following
result.

######  Theorem 4.14.

Let $C$ be a Hermitian dual-containing cyclic code over $R$ of length $n$ size
$(p^{2m})^{k}$, and let $d_{1}$ and $d_{2}$ be the minimum Hamming distances
of the $\mathrm{Res}(C)$ and $\mathrm{Tor}(C)$.Then there exists a quantum
code with parameters
$[[2n,2k-2n,\geq\mathrm{min}\\{d_{1},2d_{2}\\}]]_{p^{m}}$.

######  Example 3.

Consider cyclic codes of length $25$ over $F_{13^{2}}+uF_{13^{2}}$. In
$F_{13^{2}}+uF_{13^{2}}$,

$x^{25}-1=M_{0}(X)M_{1}(x)M_{2}(x)M_{5}(x)M_{10}(x),$

where

$M_{0}(x)=x-(1-u),$ $M_{1}(X)=x^{10}+w^{30}(1+8u)x^{5}+(1+3u),$
$M_{2}(X)=x^{10}+w^{54}(1+8u)x^{5}+(1+3u),$
$M_{5}(X)=x^{2}+w^{30}(1-u)x+(1-2u),$ $M_{10}(X)=x^{2}+w^{54}(1-u)x+(1-2u).$

Let $C=\langle M_{0}(x)M_{1}(X)^{2}M_{10}(X)^{2}\rangle$. By Theorem $4.11$,
$C^{\perp_{H}}\subset C$. Using Theorem $4.13$, we find that the Gray distance
of $C$ is equal to 4. Again from Theorem $4.14$, a $[[50,42,\geq 4]]_{13}$
quantum code may be obtained from Gray image of this code. This code is a
SQCNMDS code.

######  Example 4.

Consider cyclic codes of length $8$ over $F_{3^{4}}+uF_{3^{4}}$. In
$F_{3^{4}}+uF_{3^{4}}$,

$x^{8}-1=M_{0}(X)M_{1}(x)M_{2}(x)M_{3}(x)M_{4}(x)M_{5}(X)M_{6}(X)M_{7}(X),$

where

$M_{0}(x)=x-(1-u),~{}~{}~{}$ $M_{1}(x)=x+(1-u)w^{10},$
$M_{2}(X)=x+(1-u)w^{20},$ $M_{3}(X)=x+w^{30}(1-u),$
$M_{4}(X)=x-(1-u),~{}~{}~{}~{}$ $M_{5}(X)=x++w^{50}(1-u),$
$M_{6}(X)=x++w^{60}(1-u),$ $M_{7}(X)=x++w^{70}(1-u).$

Let $C=\langle M_{0}(x)M_{1}(X)^{2}\rangle$. By Theorem $4.11$ ,
$C^{\perp_{H}}\subset C$. Using Theorem $4.13$, we find that the Gray distance
of $C$ is equal to 3. Again from Theorem $4.14$, a $[[16,10,\geq 3]]_{9}$
quantum code may be obtained from Gray image of this code. This code is a
SQCNMDS code.

##  5 Conclusion

We give two methods to construct quantum codes from cyclic codes over finite
chain rings. Furthermore, the results show that cyclic codes over finite chain
rings are also a good resource of constructing quantum codes. We believe that
more good quantum codes can be obtained from cyclic codes over finite chain
rings. In a future work, we will use the computer algebra system MAGMA to find
more new good quantum codes.

Acknowledgements This work was supported by Research Funds of Hubei Province,
Grant No. D20144401.

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