As we briefly discussed in section 1.1, the most famous graph coloring problem is certainly the map coloring problem, proposed in the nineteenth century and finally solved in 1976.

Definition 5.8.1 A **proper coloring** of a graph
is an assignment of colors to the vertices of the graph so that no two
adjacent vertices have the same color.

Usually we drop the word "proper'' unless other types of coloring are also under discussion. Of course, the "colors'' don't have to be actual colors; they can be any distinct labels—integers, for example. If a graph is not connected, each connected component can be colored independently; except where otherwise noted, we assume graphs are connected. We also assume graphs are simple in this section.

Graph coloring has many applications in addition to its intrinsic interest.

Example 5.8.2 If the vertices of a graph represent academic classes, and two vertices are adjacent if the corresponding classes have people in common, then a coloring of the vertices can be used to schedule class meetings. Here the colors would be schedule times, such as 8MWF, 9MWF, 11TTh, etc.

Example 5.8.3 If the vertices of a graph represent radio stations, and two vertices are adjacent if the stations are close enough to interfere with each other, a coloring can be used to assign non-interfering frequencies to the stations.

Example 5.8.4 If the vertices of a graph represent traffic signals at an intersection, and two vertices are adjacent if the corresponding signals cannot be green at the same time, a coloring can be used to designate sets of signals than can be green at the same time.

Graph coloring is closely related to the concept of an
**independent set**.

Definition 5.8.5 A set $S$ of vertices in a graph is independent if no two vertices of $S$ are adjacent.

If a graph is properly colored, the vertices that are assigned a particular color form an independent set. Given a graph $G$ it is easy to find a proper coloring: give every vertex a different color. Clearly the interesting quantity is the minimum number of colors required for a coloring. It is also easy to find independent sets: just pick vertices that are mutually non-adjacent. A single vertex set, for example, is independent, and usually finding larger independent sets is easy. The interesting quantity is the maximum size of an independent set.

Definition 5.8.6 The **chromatic number** of a
graph $G$ is the minimum number of colors required in a proper coloring;
it is denoted $\chi(G)$. The
**independence number** of $G$ is
the maximum size of an independent set; it is denoted $\alpha(G)$.

The natural first question about these
**graphical parameters**
is: how small or large can they be in a graph $G$ with $n$
vertices. It is easy to see that
$$\eqalign{
1&\le \chi(G)\le n\cr
1&\le \alpha(G)\le n\cr
}$$
and that the limits are all attainable: A graph with no edges has
chromatic number 1 and independence number $n$, while a complete graph
has chromatic number $n$ and independence number 1. These inequalities
are thus not very interesting. We will see some that are more
interesting.

Another natural question: What is the relation between the chromatic number of a graph $G$ and chromatic number of a subgraph of $G$? This too is simple, but quite useful at times.

Theorem 5.8.7 If $H$ is a subgraph of $G$, $\chi(H)\le \chi(G)$.

**Proof.**

Any coloring of $G$ provides a proper coloring of $H$, simply by
assigning the same colors to vertices of $H$ that they have in
$G$. This means that $H$ can be colored with $\chi(G)$ colors, perhaps
even fewer, which is exactly what we want.

Often this fact is interesting "in reverse''. For example, if $G$ has
a subgraph $H$ that is a complete graph $K_m$, then $\chi(H)=m$ and so
$\chi(G)\ge m$. A subgraph of $G$ that is a complete graph is called a
**clique**, and there is an associated graphical
parameter.

Definition 5.8.8 The **clique number** of a graph $G$ is
the largest $m$ such that $K_m$ is a subgraph of $G$.

It is tempting to speculate that the *only* way a graph $G$
could require $m$ colors is by having such a subgraph. This is false;
graphs can have high chromatic number while having low clique number;
see figure 5.8.1. It is easy to see that this
graph has $\chi\ge 3$, because there are many 3-cliques in the
graph. In general it can be difficult to show that a graph cannot be
colored with a given number of colors, but in this case it is easy to
see that the graph cannot in fact be colored with three colors,
because so much is "forced''. Suppose the graph can be colored with 3
colors. Starting at the left
if vertex $v_1$
gets color 1, then $v_2$ and $v_3$ must be colored 2 and 3, and vertex
$v_4$ must be color 1. Continuing, $v_{10}$ must be color 1, but this is
not allowed, so $\chi>3$. On the other hand, since $v_{10}$ can be
colored 4, we see $\chi=4$.

Paul Erdős
showed in 1959 that there are graphs with arbitrarily
large chromatic number and arbitrarily large **girth** (the girth is the size of the smallest cycle in
a graph). This is much stronger than the existence of graphs with high
chromatic number and low clique number.

Bipartite graphs with at least one edge have chromatic number 2, since the two parts are each independent sets and can be colored with a single color. Conversely, if a graph can be 2-colored, it is bipartite, since all edges connect vertices of different colors. This means it is easy to identify bipartite graphs: Color any vertex with color 1; color its neighbors color 2; continuing in this way will or will not successfully color the whole graph with 2 colors. If it fails, the graph cannot be 2-colored, since all choices for vertex colors are forced.

If a graph is properly colored, then each **color
class** (a color class is the set of all vertices of a
single color) is an independent set.

Theorem 5.8.9 In any graph $G$ on $n$ vertices, $\ds {n\over \alpha}\le\chi$.

**Proof.**

Suppose $G$ is colored with $\chi$ colors. Since each color class is
independent, the size of any color class is at most $\alpha$. Let the
color classes be $V_1,V_2,\ldots,V_\chi$.
Then
$$
n=\sum_{i=1}^\chi |V_i| \le \chi\alpha,
$$
as desired.

We can improve the upper bound on $\chi(G)$ as well. In any graph $G$, $\Delta(G)$ is the maximum degree of any vertex.

**Proof.**

We show
that we can always color $G$ with $\Delta+1$ colors by a simple
**greedy algorithm**:
Pick a vertex $v_n$, and list the
vertices of $G$ as $v_1,v_2,\ldots,v_n$ so that if $i< j$,
$\d(v_i,v_n) \ge\d(v_j,v_n)$, that is, we list the vertices farthest
from $v_n$ first. We use integers $1,2,\ldots, \Delta+1$ as colors.
Color $v_1$ with 1. Then for each $v_i$ in
order, color $v_i$ with the smallest integer that does not violate the
proper coloring requirement, that is, which is different than the
colors already assigned to the neighbors of $v_i$. For $i< n$, we claim
that $v_i$ is colored with one of $1,2,\ldots,\Delta$.

This is certainly true for $v_1$. For $1< i< n$, $v_i$ has at least one neighbor that is not yet colored, namely, a vertex closer to $v_n$ on a shortest path from $v_n$ to $v_i$. Thus, the neighbors of $v_i$ use at most $\Delta-1$ colors from the colors $1,2,\ldots,\Delta$, leaving at least one color from this list available for $v_i$.

Once $v_1,\ldots,v_{n-1}$ have been colored, all neighbors of $v_n$ have been colored using the colors $1,2,\ldots,\Delta$, so color $\Delta+1$ may be used to color $v_n$.

Note that if $\d(v_n)< \Delta$, even $v_n$ may be colored with one of the colors $1,2,\ldots,\Delta$. Since the choice of $v_n$ was arbitrary, we may choose $v_n$ so that $\d(v_n)< \Delta$, unless all vertices have degree $\Delta$, that is, if $G$ is regular. Thus, we have proved somewhat more than stated, namely, that any graph $G$ that is not regular has $\chi\le\Delta$. (If instead of choosing the particular order of $v_1,\ldots,v_n$ that we used we were to list them in arbitrary order, even vertices other than $v_n$ might require use of color $\Delta+1$. This gives a slightly simpler proof of the stated theorem.) We state this as a corollary.

There are graphs for which $\chi=\Delta+1$: any cycle of odd length has $\Delta=2$ and $\chi=3$, and $K_n$ has $\Delta=n-1$ and $\chi=n$. Of course, these are regular graphs. It turns out that these are the only examples, that is, if $G$ is not an odd cycle or a complete graph, then $\chi(G)\le\Delta(G)$.

Theorem 5.8.12
**(Brooks's Theorem)**
If $G$ is a graph other than $K_n$ or $C_{2n+1}$,
$\chi\le \Delta$.

The greedy algorithm will not always color a graph with the smallest possible number of colors. Figure 5.8.2 shows a graph with chromatic number 3, but the greedy algorithm uses 4 colors if the vertices are ordered as shown.

In general, it is difficult to compute $\chi(G)$, that is, it takes a large amount of computation, but there is a simple algorithm for graph coloring that is not fast. Suppose that $v$ and $w$ are non-adjacent vertices in $G$. Denote by $G+\{v,w\}=G+e$ the graph formed by adding edge $e=\{v,w\}$ to $G$. Denote by $G/e$ the graph in which $v$ and $w$ are "identified'', that is, $v$ and $w$ are replaced by a single vertex $x$ adjacent to all neighbors of $v$ and $w$. (But note that we do not introduce multiple edges: if $u$ is adjacent to both $v$ and $w$ in $G$, there will be a single edge from $x$ to $u$ in $G/e$.)

Consider a proper coloring of $G$ in which $v$ and $w$ are different colors; then this is a proper coloring of $G+e$ as well. Also, any proper coloring of $G+e$ is a proper coloring of $G$ in which $v$ and $w$ have different colors. So a coloring of $G+e$ with the smallest possible number of colors is a best coloring of $G$ in which $v$ and $w$ have different colors, that is, $\chi(G+e)$ is the smallest number of colors needed to color $G$ so that $v$ and $w$ have different colors.

If $G$ is properly colored and $v$ and $w$ have the same color, then this gives a proper coloring of $G/e$, by coloring $x$ in $G/e$ with the same color used for $v$ and $w$ in $G$. Also, if $G/e$ is properly colored, this gives a proper coloring of $G$ in which $v$ and $w$ have the same color, namely, the color of $x$ in $G/e$. Thus, $\chi(G/e)$ is the smallest number of colors needed to properly color $G$ so that $v$ and $w$ are the same color.

The upshot of these observations is that $\ds\chi(G)=\min(\chi(G+e),\chi(G/e))$. This algorithm can be applied recursively, that is, if $G_1=G+e$ and $G_2=G/e$ then $\ds\chi(G_1)=\min(\chi(G_1+e),\chi(G_1/e))$ and $\ds\chi(G_2)=\min(\chi(G_2+e),\chi(G_2/e))$, where of course the edge $e$ is different in each graph. Continuing in this way, we can eventually compute $\chi(G)$, provided that eventually we end up with graphs that are "simple'' to color. Roughly speaking, because $G/e$ has fewer vertices, and $G+e$ has more edges, we must eventually end up with a complete graph along all branches of the computation. Whenever we encounter a complete graph $K_m$ it has chromatic number $m$, so no further computation is required along the corresponding branch. Let's make this more precise.

Theorem 5.8.13 The algorithm above correctly computes the chromatic number in a finite amount of time.

**Proof.**

Suppose that a graph $G$ has $n$ vertices and $m$ edges. The number of
pairs of non-adjacent vertices is $\na(G)={n\choose 2}-m$. The proof
is by induction on $\na$.

If $\na(G)=0$ then $G$ is a complete graph and the algorithm terminates immediately.

Now we note that $\na(G+e)< \na(G)$ and $\na(G/e)< \na(G)$: $$ \na(G+e)={n\choose 2}-(m+1)=\na(G)-1 $$ and $$ \na(G/e)={n-1\choose 2}-(m-c), $$ where $c$ is the number of neighbors that $v$ and $w$ have in common. Then $$\eqalign{ \na(G/e)&={n-1\choose 2}-m+c\cr &\le {n-1\choose 2}-m+n-2\cr &={(n-1)(n-2)\over 2}-m+n-2\cr &={n(n-1)\over 2}-{2(n-1)\over 2}-m+n-2\cr &={n\choose 2} -m -1\cr &=\na(G)-1.\cr }$$

Now if $\na(G)>0$, $G$ is not a complete graph, so there are non-adjacent vertices $v$ and $w$. By the induction hypothesis the algorithm computes $\chi(G+e)$ and $\chi(G/e)$ correctly, and finally it computes $\chi(G)$ from these in one additional step.

While this algorithm is very inefficient, it is sufficiently fast to be used on small graphs with the aid of a computer.

Example 5.8.14 We illustrate with a very simple graph:

The chromatic number of the graph at the top is $\min(3,4)=3$. (Of course, this is quite easy to see directly.)

## Exercises 5.8

**Ex 5.8.1**
Suppose $G$ has $n$ vertices and chromatic number
$k$. Prove that $G$ has at least $k\choose2$ edges.

**Ex 5.8.2**
Find the chromatic number of the graph below by using
the algorithm in this section. Draw all of the graphs $G+e$ and $G/e$
generated by the alorithm in a "tree structure'' with the complete
graphs at the bottom, label each complete graph with its chromatic
number, then propogate the values up to the original graph.

**Ex 5.8.3**
Show that $\chi(G-v)$ is either $\chi(G)$ or
$\chi(G)-1$.

**Ex 5.8.4**
Prove theorem 5.8.10
without assuming any particular
properties of the order $v_1,\ldots,v_n$.

**Ex 5.8.5**
Prove theorem 5.8.12 as follows.
By corollary 5.8.11 we need consider only regular
graphs. Regular graphs of degree 2 are easy, so we consider only
regular graphs of degree at least 3.

If $G$ is not 2-connected, show that the blocks of $G$ may colored with $\Delta(G)$ colors, and then the colorings may be altered slightly so that they combine to give a proper coloring of $G$.

If $G$ is 2-connected, show that there are vertices $u$, $v$, $w$ such that $u$ is adjacent to both $v$ and $w$, $v$ and $w$ are not adjacent, and $G-v-w$ is connected. Given such vertices, color $v$ and $w$ with color 1, then color the remaining vertices by a greedy algorithm similar to that in theorem \xrefnexternal{thm:almost brooks}{cgt.pdf}, with $u$ playing the role of $v_n$.

To show the existence of $u$, $v$, $w$ as required, let $x$ be a vertex not adjacent to all other vertices. If $G-x$ is 2-connected, let $v=x$, let $w$ be at distance 2 from $v$ (justify this), and let a path of length 2 be $v,u,w$. Use theorem 5.7.4 to show that $u$, $v$, $w$ have the required properties.

If $G-x$ is not 2-connected, let $u=x$ and let $v$ and $w$ be (carefully chosen) vertices in two different endblocks of $G-x$. Show that $u$, $v$, $w$ have the required properties.

Brooks proved the theorem in 1941; this simpler proof is due to Lovász, 1975.