%% LyX 1.1 created this file.  For more info, see http://www.lyx.org/.
%% Do not edit unless you really know what you are doing.
\documentclass[11pt]{article}
\usepackage[T1]{fontenc}
\usepackage[latin1]{inputenc}
\setlength\parskip{\medskipamount}
\setlength\parindent{0pt}

\makeatletter


%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% LyX specific LaTeX commands.
\providecommand{\LyX}{L\kern-.1667em\lower.25em\hbox{Y}\kern-.125emX\@}

%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% User specified LaTeX commands.
\usepackage[T1]{fontenc}
\usepackage[latin1]{inputenc}
\setlength\parskip{\medskipamount}
\setlength\parindent{0pt}
\usepackage{graphics}

\makeatletter


\usepackage[T1]{fontenc}
\usepackage[latin1]{inputenc}
\setlength\parskip{\medskipamount}
\setlength\parindent{0pt}

\makeatletter


\setlength{\oddsidemargin}{0in}
\setlength{\evensidemargin}{0in}
\setlength{\topmargin}{0.0in}
\setlength{\headheight}{0in}
\setlength{\headsep}{0.5in}
\setlength{\textwidth}{6.5in}
\setlength{\textheight}{8.5in}
\setlength{\parindent}{0in}
\setlength{\parskip}{0.12in}

\usepackage{times}
\usepackage{psfig}
\usepackage{fancyheadings}


\makeatother


\makeatother


\makeatother

\begin{document}

\pagestyle{fancy} \cfoot{CRC: \thepage}


\title{Cyclic Redundancy Checking}


\author{Norman Matloff \\
 Dept. of Computer Science \\
 University of California at Davis\\
  \copyright{}2001, N. Matloff}


\date{May 8, 2001}

\maketitle

\section{The Importance of Error Detection}

A transmitted bit can be received in error, due to {}``noise{}'' on the transmission
channel. If we are dealing with voice or video data, the occurrence of errors
in a small percentage of bits is quite tolerable, but in many other cases it
is crucial that all bits be received intact. If for example we are downloading
a binary program file, the program may be unexecutable if even one bit is incorrect.
If the destination or source address of a packet has one bit wrong, communication
is impossible.

There are many methods which have been developed to detect errors, applied at
different levels of the seven-layer model. Of course, no method can detect \underbar{all}
errors, but a number of methods in use today are amazingly effective. The one
we will discuss here is the famous CRC coding form.


\section{The CRC Error-Detection Method}

One powerful error-detection method is Cyclic Redundancy Checking (CRC), a generalization
of parity checking. One appends a few (typically 16 or 32) bits to the end of
the bit string for a message and sends out the extended string. The receiver
then performs a computation which would yield 0 if no bits of the message had
been in error; if the result is not 0, then the receiver knows that there has
been an error in one or more bits.

CRC cannot detect all possible errors, but the range of errors which it \textbf{can}
detect is impressively broad. As networks textbook authors Larry Peterson and
Bruce Davie have pointed out, it is quite remarkable, for instance, that a mere
32 bits of CRC are sufficient to provide adequate protection against errors
in the 12,000-bit messages which Ethernet can send.


\section{How CRC Works}

Let M be the message we wish to send, m bits long. Let C be a divisor string,
c bits long. C will be fixed (say hardwired into a serial I/O chip), while M
(and m) will vary. We must have m \( > \) c-1. It is always assumed that c
\( > \) 1 and that the first and last bits of C are 1s.

Our CRC field will consist of a string R, c-1 bits long. Here is how to generate
R, and send the message:

\begin{itemize}
\item 1. Append c-1 0s to the right end of M. These are placeholders for where the
CRC will go. Call this extended string M'. 
\item 2. Divide M' by C, using mod-2 arithmetic. \textit{Note carefully: This is mod-2,
not base-2. These quantities are NOT numbers in the usual sense.}


Note that each time we divide C, which is a c-bit quantity, into another c-bit
quantity Z, we say that C {}``goes into{}'' Z 1 time if the leading bit of
Z is 1, or 0 times otherwise.

Call the remainder R. Since we are dividing by a c-bit quantity, R will be c-1
bits long.

\item 3. Replace the c-1 appended 0s in M' by R. Call this new string W. 
\item 4. Send W to the receiver. 
\end{itemize}
For example, say C = 1011 and M = 1001101. Then we divide as follows:

%% adapted from Jeremy Gibbons, TUGboat, 18 (1997), No.2
\newdimen\digitwidth
   \settowidth\digitwidth{0}
\def~{\hspace{\digitwidth}}
\def\divrule#1#2{%
   \noalign{\moveright#1\digitwidth%
      \vbox{\hrule width#2\digitwidth}}}
1011\,\begin{tabular}[b]{@{}r@{}}
        1010011 \\ \hline
        \big)\begin{tabular}[t]{@{}l@{}}
        1001101000 \\
        1011 \\ \divrule{0}{4}
        ~~1010 \\
        ~~1011 \\ \divrule{2}{4}
        ~~~~~1100 \\
        ~~~~~1011 \\ \divrule{4}{4}
        ~~~~~~1110 \\
        ~~~~~~1011 \\ \divrule{6}{4}
        ~~~~~~~101
   \end{tabular}
\end{tabular}

We first divide 1011 into 1001 (see above), with a quotient of 1. Yes, that's
right! As long as the dividend has at least as many digits as the divisor--both
have 4 digits here--then the divisor does indeed {}``go into{}'' the dividend
with a quotient of at least 1. If they have the same number of digits,\footnote{%
The number of digits is counted beginning with the first 1 digit, so for example
0011 is considered only a 2-digit quantity, not a 4-digit one. 
} the quotient is 1, which is placed as the first digit in the overall quotient
above the division sign. Our remainder is 0010 (remember, our arithmetic here
is mod 2, bitwise), which we write as 10.

Next we extend that 10 by bringing down the next digit in the dividend, a 1,
forming 101, that is 0101. 1011 goes into 0101 0 times, so we place this 0 as
the next digit in the overall quotient above, and bring down the next digit
of the dividend, a 0, making 1010. Now 1011 goes into 1010 1 time, and so on.

The final remainder is 101, so the transmitter now sends the string 1001101101.

(You should check that one can {}``multiply{}'' back to get the original result,
i.e. 1011{*}1010011 + 101 = 1001101000. Remember, though, that the multiplication
and addition are done on a mod-2 basis, i.e. XOR with no carries.)

The receiver then receives a string Y, which is hopefully the same as W but
might not be, due to line noise. The receiver divides Y by C; if the remainder
is nonzero, the receiver decides Y had one or more bits in error, while if the
remainder is zero, the receiver accepts Y as correct (though it still might
not be).


\section{Why It Works}

Note that the replace operation in step 3 above is equivalent to performing

\begin{equation}
W\leftarrow M'+R
\end{equation}


Moreover, because we are using mod 2 arithmetic, that operation is also equivalent
to

\begin{equation}
W\leftarrow M'-R
\end{equation}


If one divides, say, 63 by 5, then one gets a remainder of 3. That means that
if we subtract that remainder of 3 from the dividend 63, the result 60 will
be evenly divisible by 5.

Similarly, since R was our remainder from M'/C, we see from the last equation
that M'-R is exactly divisible by C. So, W is divisible by C.

Let E denote the \textbf{error vector}, a symbolic way to describe which bits
in W get corrupted on their way to the receiver. If for instance the i-th bit
of E is a 1, this means that there was an error in the i-th bit of W. Then the
receiver will receive Y = W + E, again keeping in mind that this is mod-2, bitwise
addition.

We hope E consists of all 0s, of course, but it might not. The receiver will
divide Y by C; since W is evenly divisible by C, then Y will be too if E = 0.
So, if the receiver finds that Y/C has a nonzero remainder, the receiver can
be sure that there was an error.

In other words, the CRC will detect all error patterns except those for which
E is evenly divisible by C. Clever people have found good strings C for which
a remarkably small proportion of Es are evenly divisible by C. These Cs have
been publicized and are in common use.

Note that E is an m+(c-1) bit quantity, so it includes the bits corresponding
to the R portion of W. So, the system here is even capable of detecting some
errors in the CRC field itself!

By the way, when we divide M' by C, getting a quotient Q and a remainder R,
Q will be m bits long. To see this, draw a {}``long division{}'' diagram as
above, and note that the first and last bits of Q will be above the c-th and
(m+c-1)-st bits of M', so that Q consists of (m+c-1) - c + 1 = m bits.


\section{What Kinds of Errors Will CRC Detect?}

In CRC analysis, it helps to represent bit strings as polynomials. For instance,
our C above, 1011, would be denoted by

\begin{equation}
\label{cpoly}
x^{3}+x+1
\end{equation}


The word {}``denoted{}'' here is key; the polynomial is just notation, and
there is no physical variable x. But this notation turns out to be an extremely
powerful tool for analyzing CRC. We will use C(x), E(x) and so on to refer to
the polynomial versions of the bit strings C, E etc.

The question at hand is which types of E strings are detectable by which C strings.
Remember, we decide what string to use for C, say when we design a serial I/O
chip or Ethernet interface card. So, we want to choose C in such a way that
many classes of E strings are detectable. Recall that an E string will be detected
as long as C does not evenly divide it, and our analysis will stem from this
consideration.

Remember, any C is eligible for use in CRC as long as its first and last bits
are 1s, and as long as \( c\geq 2 \) That means that the corresponding polynomial,
C(x), has \( x^{c-1} \) as its leading term, and 1 as its last term. This is
illustrated in Equation (\ref{cpoly}) in the case of C = 1011, c = 4.

Below are two examples of the many properties CRC theorists have proven hold
for any CRC-eligible C. Before presenting those proofs, it would be useful to
review some basic facts about prime factorization. Let's abandon the odd arithmetic
in CRC for a moment, and consider ordinary integers u, v and w, where w = uv.

As you know, each integer has a prime factorization, e.g. \( 120=2\times 2\times 2\times 3\times 5 \).
Thus, each of u , v and w has a prime factorization, say \( u=p_{u1}p_{u2}...p_{ur} \),
\( v=p_{v1}p_{v2}...p_{vs}, \) and \( w=p_{w1}p_{w2}...p_{wt} \). But since
w = uv, we know that a more detailed representation of the prime factorization
of w is \( p_{u1}...p_{ur}p_{v1}...p_{vs} \). In other words:

\textbf{If w = uv, then any prime factor of w must be a prime factor of either
u or v. }

This simple fact will be useful in our proofs below, even though the entities
we are working with are now polynomials rather than integers. 

Our proofs below will be by contradiction. We are trying to show that C(x) does
not divide E(x), but will suppose that it does, and run into a contradiction.

Here is a sampling of what can be proved about any CRC-eligible C:

\begin{itemize}
\item \textbf{Any single-bit error will be detected.}


To see this, note that for such an error pattern, E will consist of all 0s except
for a single 1 bit in some position within the string. So, E(x) will be \( x^{k} \)
for some k. Note that this is already in prime-factorization form, x being the
prime factor (i.e. it cannot be broken down further into other factors), repeated
k times.

First consider the case in which k > 0. Suppose C(x) were to evenly divide \( x^{k} \),
i.e. \( x^{k}=C(x)D(x) \) for some polynomial D(x). That would imply that x
is a factor of C(x) {[}and of D(x){]}, from the discussion on prime factorization
above. Yet x could not be a factor of C(x) due to the presence of the constant
term 1 in C(x). (Recall that C(x) is required to have the form \( x^{c-1}+...+1 \))

In the case k = 0, i.e. E(x) = 1, then the fact that C(x) has at least one other
term than 1 (namely \( x^{c-1} \)) means C(x) could not divide E(x). 

So, this E(x) is guaranteed not to be evenly divisible by C(x), and thus a single-bit
error will be caught by the CRC.

\item \textbf{Any burst error, i.e. a set of consecutive bit errors, of length at
most c-1 will be detected.}


In our worked-out example above in which W = 1001101101, if E is, say 0000111000,
with consecutive errors in the fifth, sixth and seventh bits of W, then such
a pattern of errors will be detected.

To prove this, let w denote the length of W, which is also the length of E.
Say the most significant 1 in E is in bit position j, where the least significant
position is called position 0 and the most significant position is numbered
w-1. Denote the length of the burst by b. Then E(x) will have the form

\begin{equation}
x^{j}+x^{j-1}+x^{j-2}+...+x^{j-b+1}
\end{equation}


which can be factored as

\begin{equation}
x^{j-b+1}(x^{b-1}+...+1)
\end{equation}


Suppose C(x) were to evenly divide this E(x), i.e. E(x) = C(x) D(x) for some
polynomial D(x). Note again that each x factor in \( x^{j-b+1} \) is prime,
and as in our previous discussion x is NOT a prime factor of C(x). So, all of
C(x)'s prime factors would have to be among the prime factors of \( x^{b-1}+...+1 \).
In other words, C(x) would have to evenly divide \( x^{b-1}+...+1 \) Yet that
is impossible, since b was given to be at most c-1 and thus \( x^{b-1}+...+1 \)
has degree less than that of C(x) (the former has degree b-1, the latter degree
c-1).

So we see C(x) cannot evenly divide E(x), and our claim is proved.

This property of catching burst errors is quite important, as it is often the
case that line noise will indeed occur in bursts.

\end{itemize}

\section{Digital Implementation}

As you can see, the CRC algorithms could easily be encoded in software, but
that may be too slow. It would be much faster if implemented in digital logic
hardware, and it turns out that this can be done quite simply.

As before, let c denote the number of bits in the divisor polynomial C(x), so
that the degree of the polynomial is c-1. Then we set up a left-to-right shift
register with c-1 bits, which we will denote, from left to right, \( S_{0},S_{1},...,S_{c-2}. \).
As with any shift register, at each clock impulse \( S_{i} \) is replaced by
the value which had been in \( S_{i-1} \), where we are calling the left input
{}``\( S_{-1} \).{}''

However, our shift register will differ from an ordinary one in that at each
clock pulse the value which had been in \( S_{c-1} \) will be recycled and
XOR-ed with some of the other \( S_{i} \). Specifically, if there is a term
\( x^{j} \) in the divisor polynomia C(x), then there is an XOR just to the
left of \( S_{i} \) (except for i = c-1).

The message M, with c-1 0s appended as before, is fed into the input to the
shift register, one bit per clock cycle, starting with M's most significant
(i.e. leftmost) bit. After m+c clock cycles, the shift register contains the
remainder (i.e. the CRC field), in reverse order.

Our figure here shows the setup for the case \( C(x)=x^{3}+x+1 \), i.e. C =
1011, from our example above.

\vspace{0.3cm} \par\centering \includegraphics{DigCRC.eps} \par{} \vspace{0.3cm}

The table here shows a pulse-by-pulse account of what will occur with the input
M = 1001101 (with three 0s appended):

\vspace{0.3cm} \centering 

\begin{tabular}{@{}l@{}}
clk\\
 input\\
 S0\\
 S1\\
 S2\\
0\\
 1\\
 0\\
\end{tabular}

\par{} \vspace{0.3cm}

For example, at the end of clock period 3, the register contains (0,0,1), with
an input of 1 coming in, say, near the end of the clock period, taking effect
in the next period. Then the clock pulse starting period 4 comes. What happens?

\begin{itemize}
\item The new \( S_{0} \) will be the XOR of 1 (from the old \( S_{2} \)) and 1
(from the input shown in row 3), which is 0. Thus the new \( S_{0} \) will
be 0. 
\item The new \( S_{1} \) will be the XOR of 1 (from the old \( S_{2} \)) and 0
(from the old \( S_{1} \)), i.e. 1. 
\item The new \( S_{2} \) will be equal to the old \( S_{1} \), i.e. 0 
\end{itemize}
So, the new register state will be (0,1,0), seen in row 4.

This would occur at the transmitter end, and a similar setup for the CRC check
will be at the receiver end.
\end{document}