A linear algorithm for the grundy number of a tree

International Journal of Computer Science & Information Technology (IJCSIT) Vol 6, No 1, February 2014
DOI : 10.5121/ijcsit.2014.6112 155
A Linear Algorithm for the Grundy Number of A
Tree
Ali Mansouri and Mohamed Salim Bouhlel
Department of Electronic Technologies of Information and Telecommunications Sfax,
Tunisia
ABSTRACT
A coloring of a graph G = (V ,E) is a partition {V1, V2, . . . , Vk} of V into independent sets or color classes.
A vertex v ∈Vi is a Grundy vertex if it is adjacent to at least one vertex in each color class Vj for every j <i.
A coloring is a Grundy coloring if every color class contains at least one Grundy vertex, and the Grundy
number of a graph is the maximum number of colors in a Grundy coloring. We derive a natural upper
bound on this parameter and show that graphs with sufficiently large girth achieve equality in the bound.
In particular, this gives a linear-time algorithm to determine the Grundy number of a tree.
KEYWORDS
Graphs, Grundy Number, Algorithm , Linear-time algorithm , Coloring of a graph.
1. INTRODUCTION
In this paper, we consider the calculation of the Grundy number of a graph with large girth. The
concept of Grundy coloring is a weak form of minimal coloring.
Let G = (V ,E) be a graph. A set of vertices S ⊆V is called independent if no two vertices in S
are adjacent. A (proper) k-coloring of graph G is a partition {V1, V2, . . . , Vk} of V into k
independent sets, called color classes. In the coloring, a vertex v ∈Vi is called a Grundy vertex if
v is adjacent to at least one vertex in color class Vj , for every j <i. (Every vertex in V1 is
vacuously a Grundy vertex.) Thus the Grundy (coloring) number of a graph is defined as the
maximum number of colors in a coloring where every vertex is a Grundy vertex. This parameter
is studied in [1–3,5–8,11] inter alia.
Recently, Erdös et al. [4] studied the concept of a partial Grundy coloring. This is a coloring (V1,
V2, . . . , Vk) such that every color class Vi, 1<i<k, contains at least one Grundy vertex. The
Grundy number Ə Γ (G) is the maximum k such that G has a k-Grundy coloring [4]. Erdös et al.
related Grundy colorings to other graph properties such as parsimonious proper coloring number
[10] and the maximum degree.
In this paper, we observe a natural upper bound on Ə Γ (G) and show that trees and graphs with
sufficiently large girth achieve equality in the bound. In particular, this gives a linear time
algorithm to determine the Grundy number of a tree. As expected, the associated decision
problem is NP-complete

156
2. AN UPPER BOUND
In this section, we define a natural upper bound on the Grundy number in terms of what we call a
“feasible Grundy sequence”. We show that a greedy algorithm can compute this bound on all
graphs.
Definition. A sequence S of r distinct vertices (g1, . . . , gr ) of a graph G is a feasible Grundy
sequence if for 1<i< r the degree of gi in G − {g i+1, . . . , gr } is at least i − 1. The stair factor ‫ל‬ (G)
of G is the maximum cardinality of a feasible Grundy sequence.
For example in the graph in Fig. 1, the sequence (b, c, d, e) is a feasible Grundy sequence.
Note that the stair factor is clearly no greater than the maximum degree plus 1. If we have a
Grundy coloring and write down a Grundy vertex from each color class from Vk down to V1, we
obtain a feasible Grundy sequence. Thus it follows that:
Ə Γ (G) < ‫ל‬ (G) < ∆ (G) + 1.
Surprisingly perhaps, ‫ל‬ (G) can be calculated in linear-time with a simple greedy algorithm. (This
has echoes in the calculation of the so-called coloring number of a graph; see [9]). We introduce
the following lemma.
Lemma 2.1. Let w be a vertex of graph G. Then ‫ל‬ (G − w)> ‫ל‬ (G) − 1.
Proof. Consider a maximum feasible Grundy sequence S of graph G: say (g1, . . . , g‫ל‬ ). If w ∈S,
then S − w is a feasible Grundy sequence in graph G− w. If w /∈S, then S − g1 is a feasible
Grundy set in G − w. Hence the result.
Figure 1. An example of feasible Grundy sequence.
We next introduce a vertex-decomposition list.
Definition. A vertex-decomposition list D = (v1, . . . , vn) of a graph is computed by repeatedly
removing a vertex vi of maximum degree from the graph Gi =G−{v1, . . . , v i−1}. (Note that v1 is a
maximum degree vertex in G.) The degree of the vertex vi at removal we call its residue degree,
and denote the sequence of residue degrees by (d1, . . . , dn).
Theorem 1. Let D be any vertex-decomposition list of a graph G with residue degrees (d1, . . . ,
dn). Then
‫ל‬ (G) = min di + i.

157
Proof. Let k be the value mini di + i. It is clear that (v1, . . . , vk) is a feasible Grundy sequence.
Hence ‫ל‬ (G )> k.
We prove that ‫ל‬ (G) < k by induction on the order n. If n=1, then the graph is an isolated vertex,
and so ‫ל‬ (G) = k = 1. In general, let G’= G − v1 and let k’ be the value produced by the algorithm
on G’ using D’ = D − v1. That is, k’ = mini> 2 di + (i − 1).
There are two cases.
Case 1: k’< k. Then ‫ל‬ (G) < k’ by the inductive hypothesis, and so ‫ל‬ (G) < ‫ל‬ (G’) + 1 = k’ + 1< k
by Lemma 2.1.
Case 2: k’ = k. Then the only way that mini >2 di + i − 1 = mini> 1 di + i = k is that d1 = k − 1.
That is, ∆ (G) = k − 1. But clearly ‫ל‬ (G) < ∆ (G) + 1 in all graphs.
2.1. Complexity analysis
We show that the parameter ‫ל‬ (G) can be computed in O (m) time, where m is the number of
edges in the graph. We assume the graph is given by its adjacency list representation.
We proceed as follows. We use bucket sort to distribute the n vertices into an array of n buckets
corresponding to residue degree 0 through n − 1. While calculating the decomposition sequence,
we find the vertex vi in the last non-empty bucket and remove it. Then for each entry u in vi ’s
adjacency list, we move u one bucket down.
To allow each bucket move to take O (1) time, we store each bucket as a doubly linked list; and
also maintain for each vertex vi a pointer that tracks the location of its entry in the bucket-array.
During a run of the algorithm, the total number of moves is at most the sum of the degrees and is
hence O (m). Therefore the algorithm has complexity O (m).
3. ACHIEVING GRUNDY COLORING
In this section we show that, in graphs with sufficiently large girth, any feasible Grundy sequence
is always “almost realizable” in a Grundy coloring, and present a linear time algorithm for
achieving this. This implies that on such graphs the partial Grundy number is equal to the upper
bound. We prove the following theorem.
Theorem 2.
Let G be a graph with girth larger than 8 and let S=(g1, . . . , gk) be a feasible Grundy sequence
for G. Then there exists a partial Grundy coloring (V1, . . . , Vl) with l > k such that for each i >2,
gi is in Vi and is a Grundy vertex.
Note that the restriction i>2 is necessary: consider for example a path on 4 vertices labeled g1–g3–
v–g2. In this case, the vertex labeled g1 is given color 2 following
Theorem 2.
We first give the following definition.
Definition. Consider a feasible Grundy sequence (g1, . . . , gk). Vertex gi is satisfied if we properly
color some of its neighbors in such a way that gi is adjacent to colors {1, . . . , i−1},but i is not

158
among the colors now appearing on its neighbors. (In other words, gi becomes a Grundy vertex
for color class Vi)
The basic approach in proving Theorem 2 is as follows:
We color the graph in k −1 rounds. In Round i (i running from 2 up to k), we color the vertex gi
with color i and then satisfy it. At the end, we color all the uncolored vertices (if any) one at a
time by giving each the smallest color not used in its neighborhood.
To satisfy gi , we must color i−1 of its neighbors with distinct colors. So we must ensure that each
round of coloring does not create problems for later rounds: specifically it must ensure that
already-colored neighbors have distinct colors. For this we introduce the concept of a restriction.
Definition. Let w be an uncolored neighbor of gi . Then a restricted color for w, is any color at a
neighbor of w, and any color at distance 2 from w, where the common neighbor of w and the
colored vertex is one of g i+1, . . . , gk.
Note that at the start of round i, every colored vertex is either one of g2, . . . , g i−1, or has been
colored to satisfy one of these vertices (its coloring was caused by that Grundy vertex). Thus
there are 4 cases of restricted colors,. (Note that in case (1), w was not used to satisfy its adjacent
Grundy vertex.) The key is that the girth limits the number of restrictions. Note that when
restrictions are not violated, no two vertices in Ni (see below) are already colored with the same
color at round i.
Lemma 3.1. Let graph G have girth larger than 8. Let Ni be any set of i − 1 neighbors of gi
excluding g i+1, . . . , gk. If at the start of round i no two vertices in Ni are already colored with the
same color, then we can complete the coloring of Ni such that every vertex in Ni has a distinct
color in the range {1, . . . , i − 1} and the color restrictions are not violated.
Proof. Assume that no restrictions have been violated up to the start of round i. Before round i,
only colors 1 up to i − 1 have been used. Consider a vertex g j for 2< j <i. Then, by the girth
constraint, either gj caused one colored neighbor of gi and no restriction on the uncolored vertices
of Ni, or gj caused no colored neighbor and at most one restriction.
Hence if b of the elements of Ni are uncolored, then there are at most b−1 restrictions and b
available colors. It follows that we can complete the coloring of Ni greedily—provided we color
the vertices that have a restriction first; each time, we remove the color used from the available
list at each vertex. Because the restrictions are not violated and available list is updated after each
coloring, no two vertices in Ni are colored with the same color while satisfying gi . No restrictions
have been violated in the execution of round i.
We observe that a set N always exists. So this completes the proof of Theorem 2. As a
consequence, we obtain:
Theorem 3. For a graph G of girth larger than 8, ‫ל‬(G) = Ə Γ (G).
4. THE GRUNDY (COLORING) NUMBER
As expected, the Grundy number is intractable. Define the following decision problem:
GRUNDY COLORING:
Instance: Graph G, positive integer k
Question: Does G have a partial Grundy coloring ∏ = {V1, V2, . . . , Vk} with at least k colors?

159
Theorem 4. GRUNDY COLORING is NP-complete, even for chordal graphs.
Proof. Clearly GRUNDY COLORING is in NP. The certificate is the coloring.
The transformation is from 3-COLOR (is a graph 3-colorable?). Given an arbitrary instance G of
3-COLOR, we create a graph G’ and integer k, such that G has a (proper) 3-coloring if G’ has a
Grundy k-coloring for G’.
The graph G’ is constructed as follows:
1. For each vertex vi ∈V (G) create a single vertex labeled vi in G’. For each edge ej ∈E (G)
create a single vertex labeled lj in G’. Let R be the set of these vertices (both vi and lj) in G’.
2. For each edge ej =vavb ∈E (G), create a vertex sj in G’, and add edges from sj to va,vb and lj in
G’. Then form all the sj vertices into a clique S in G’.
3. Finally, add one disjoint K3 in G’.
An example of this construction is shown in Fig. 2.
G G’
Figure 2. An example
Clearly, this construction is polynomial in the size of G. Let k = |E| + 3 (the maximum degree of
G’ plus 1).
First we show that if G has a 3-coloring g, then G’ has a Grundy coloring f that uses k colors. We
construct f as follows:
1. For vi ∈V l et f (vi )=g(vi ). For each lj in G’ representing an edge ej =vavb in G, let f (lj ) be
whichever of 1, 2 or 3 is not used for va and vb.
2. For j = 1, . . . , |E|, let f (sj ) = 3 + j (in any order).
3. Color the K3 with 1, 2 and 3.
The assignment is a valid partial Grundy k-coloring. For, each color class Vi is an independent
set. Note that for i > 4, each color class Vi has only one member, namely s i−3.
Each sj vertex is adjacent to vertices colored 1, 2 and 3 in R (va, vb, and lj ) and to vertices colored
4, . . . , j − 1.
Now we show that if G’ has a Grundy coloring f using k colors, then G has a proper 3-coloring.
Indeed, we claim that the restriction of f to V is a proper 3-coloring of G. Note the following:
1. Suppose some vertex of R is Grundy in G’. Let w be such a vertex with the largest color c, c>1.
Vertex w is only adjacent to vertices of S; hence S must contain vertices of colors 1 through c − 1.
By the choice of w, S contains Grundy vertices for colors c + 1 through k. This implies that |S| >

160
k − 1, a contradiction. Therefore, in graph G’, no vertex of R is Grundy with color larger than 1.
[12]
2. Thus S consists of the Grundy vertices for colors 4 up to k. Each of these vertices has three
neighbors in R: these must be colored with colors 1, 2 and 3, and must have distinct colors.
3. Thus V is colored with only colors 1, 2 and 3. Furthermore, if va and vb are adjacent in G, there
is a vertex of S adjacent to both of them and so they receive different colors. That is, f restricted
to V is a proper coloring of G.
A similar construction shows that the partial Grundy number is intractable even for bipartite
graphs.
4. CONCLUSION
In this paper we show that the problem of determining the Grundy number Ə Γ (G) of a graph G
is NP-complete even for chordal graphs and bipartite graphs. We also show that Ə Γ (G) for any
G is bounded above by the stair factor ‫ל‬ (G), a parameter newly defined here.
For graphs with girth larger than 8, this bound is achieved. This result leads to a linear time
algorithm to determine Ə Γ (G) for graphs with large girth.
We also observe that Ə Γ (G) <‫ל‬ (G) <∆ (G) + 1 for any graph G. This inequality chain raises
some interesting questions. We know from Theorem 3 that for a graph G with girth larger than 8,
Ə Γ (G) = ‫ל‬ (G).
ACKNOWLEDGEMENTS
I think Mr. Zhengnan Shia for his collaboration.
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