Semantical evaluations as monadic second-order

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Graph structure and
Monadic second-order logic
Bruno Courcelle
Université Bordeaux 1, LaBRI
References :
Chapter 5 in : Handbook of graph grammars vol.1, 1997,
Book in progress,
Articles with J. Makowsky, U. Rotics, P. Weil, S. Oum, A. Blumensath
See :
Graph structure :
Embedding in a surface
Exclusion of minor, vertex-minor, induced subgraph
Monadic second-order logic : expression of properties, queries, optimization
functions, number of configurations.
Mainly useful for tree-structured graphs (Second-order logic useless)
Two types of questions :
Checking G =  for fixed MS formula , given G
Deciding  G, G  C, G =  for fixed C, given MS formula .
Tools to be presented
Algebraic setting for tree-structuring of graphs
Recognizability = finite congruence  inductive computability
 finite deterministic automaton on terms
Fefermann-Vaught : MS definability  recognizability.
History : Confluence of 4 independent research directions,
related :
1. Polynomial algorithms for NP-complete and other hard problems on particular
classes of graphs, and especially hierarchically structured ones : series-parallel
graphs, cographs, partial k-trees, graphs or hypergraphs of tree-width < k, graphs of
clique-width < k.
2. Excluded minors and related notions of forbidden configurations (matroid
minors, « vertex-minors »).
3. Decidability of Monadic Second-Order logic on classes of finite graphs, and on
infinite graphs.
4. Extension to graphs and hypergraphs of the main concepts of Formal
Language Theory : grammars, recognizability, transductions, decidability questions.
1. Introduction
Extension of Formal Language Theory notions
2. Recognizability, an algebraic notion.
3. Context-free sets defined by equation systems.
4. The graph algebras VR and HR.
Algorithmic applications :
5. Inductive computations and recognizability; fixed-parameter tractable algorithms.
6. Monadic second-order logic defines inductive properties and functions
Formal language theory extended to graphs
7. Closure and decidability properties ; generation of classes of graphs by monadic
second-order transductions.
Monadic second-order logic and combinatorics
Seese's conjecture
Open questions
Introduction : An overview chart :
sets of graphs
Language theory
for graphs
Monadic 2nd
sets of graphs
Mon. 2nd order transductions
order logic
Fixed parameter
tractable algorithms
Key concepts of FLT and their extensions
Algebraic structure :
monoid (X*,*,)
Algebras based on graph operations :  //
quantifier-free definable operations
Algebras : HR, VR
Equational sets of the
algebras HR,VR
Context-free languages :
Equational subsets of (X*,*,)
Regular languages :
Finite automata 
Finite congruences
Regular expressions
Monadic Second-order
Recognizable sets
of the algebras
defined by congruences
definable sets of words or terms Monadic Second-order definable sets of graphs
Rational and other types of
Monadic Second-order transductions
Relationships between algebraic and logical notions
for sets of graphs (and hypergraphs)
notions characterizations characterizations
union, Rec
equation systems MS-trans(Trees)
Boolean opns
MS-def REC
Signatures for graph algebras :
HR :
graphs and hypergraphs with “sources”
VR :
graphs with vertex labels, “ports”
VR+ : VR with quantifier-free operations (ex. edge complement)
Another picture :
Value ( MS Transduction)
MS Transductions
(MS Transductions)
MS Transduction
Binary trees
Equational sets = MS-Trans(Binary Trees)
Context-free languages = images of the Dyck language
(which encodes trees) under rational transductions
Since MS transductions are closed under composition, the
class of equational sets is closed under MS transductions
2. Recognizable sets :
algebraic definition
F : a finite set of operations with (fixed) arity.
M = < M, (fM)f  F > : an F-algebra.
Definition : L
 M is (F-)recognizable if it is a union of equivalence classes for a finite
congruence  on
M (finite means that M /  is finite).
Equivalently, L = h-1(D) for a homomorphism h : M
 A, where A is a finite F-
algebra, D  A. (On terms : Finite deterministic automata).
REC(M) is the set of recognizable subsets of M, with respect to the algebra M.
Closure properties : REC(M) contains M and , and is closed under union, intersection
and difference.
The many-sorted case with infinitely many sorts
S : the countable set of sorts.
F : an S-signature : each f in F has a type s1s2 …sk  s, with s, si  S
M = < (Ms)s  S, (fM)f  F > F-algebra, Ms Mt = , if s t
where fM : Ms1 X Ms2 X … X Msk  Ms
Definition : L
 Ms is
(F-) recognizable if it is a union of equivalence classes for a
congruence  on M such that equivalent elements are of the same sort and there are
finitely many classes of each sort.
3. Equational (context-free) sets
Equation systems = Context-Free (Graph) Grammars
in an algebraic setting
In the case of words, the set of context-free rules
S aST;
S b ; T cTTT; T a
is equivalent to the system of two set equations:
S = aST
T = cTTT
where S is the language generated by S
(idem for T and T).
For graphs (or other objects) we consider systems of equations like:
S = f( k( S ), T )
 {b}
T = f( T , f( g(T ), m( T )))  { a }
where f is a binary operation, g, k, m are unary operations on graphs,
a, b denote basic graphs (up to isomorphism).
An equational set is a component of the least (unique) solution of such an
equation system. This is well-defined in any algebra.
Closure properties and algebraic characterizations
General algebraic properties
notions characterizations
EQ equation systems
union, Rec
Val(REC(Terms)) homomorphisms
Boolean operations
Theorem (Mezei and Wright) :
1) In an algebra of terms T(F) : EQ(T(F)) = REC(T(F))
2) In an F-algebra M :
where ValM : T(F)
EQ(M) = ValM(REC(T(F))
M is the evaluation mapping, the unique homomorphism.
4. The graph algebras VR and
Tree-width : Tree-decomposition of width k : k+1 = maximum size of a box
Tree-width : twd(G) = minimum width of a tree-decomposition
Trees have tree-width 1, Kn has tree-width n-1, the n x n grid has tree-width n
Outerplanar graphs have tree-width at most 2.
HR operations : Origin :
Hyperedge Replacement hypergraph grammars ; associated complexity
measure : tree-width
Graphs have distinguished vertices called sources, pointed to by labels from a set of
size k :
{a, b, c, ..., h}.
Binary operation(s) :
G // H
Parallel composition
the disjoint union of G and H and sources with same label are fused. (If
G and H are not disjoint, one first makes a copy of H disjoint from G).
Unary operations :
Forget a source label
without a-source : the source is no longer
distinguished ; it is made "internal".
Source renaming :
Rena,b(G) exchanges source names a and b
(replaces a by b if b is not the name of a source)
Nullary operations denote basic graphs : the connected graphs with at most one edge. For
dealing with hypergraphs one takes more nullary symbols for denoting hyperedges.
More precise algebraic framework : a many sorted algebra where each finite set of
source labels is a sort. The above operations are overloaded.
Proposition: A graph has tree-width  k if and only if it can be constructed
basic graphs with  k+1 labels by using the operations
// , Rena,b and Forgeta .
Example : Trees are of tree-width 1, constructed with two source labels, r (root) and n (new root):
Fusion of two trees at their roots :
Extension of a tree by parallel composition
with a new edge, forgetting the old root,
making the "new root" as current root :
E = r  n
Renn,r (Forgetr (G // E))
From an algebraic expression to a tree-decomposition
Example : cd // Rena,c (ab // Forgetb(ab // bc))
Constant ab denotes a directed edge from a to b.
The tree-decomposition associated with this term.
Origin : Vertex Replacement graph grammars
Associated complexity measure :
clique-width, has no combinatorial characterization but is
defined in terms of few very simple graph operations (whence easy inductive proofs).
Equivalent notion : rank-width (Oum and Seymour) with better structural and algorithmic
Graphs are simple, directed or not.
k labels : a , b , c, ..., h. Each vertex has one and only one label ;
a label p may label several vertices, called the p-ports.
One binary operation: disjoint union
: 
Unary operations: Edge addition denoted by Add-edga,b
Add-edga,b(G) is G augmented with (un)directed edges from every
to every
Vertex relabellings :
Relaba,b(G) is G with every vertex labelled by a relabelled into b
Basic graphs are those with a single vertex.
Definition: A graph G has clique-width  k it can be constructed from basic graphs
by means of k labels and the operations , Add-edga,b and Relaba,b
Its (exact) clique-width, cwd(G), is the smallest such k.
Proposition : (1) If a set of simple graphs has bounded tree-width, it has bounded
clique-width, but not vice-versa.
(2) Unlike tree-width, clique-width is sensible to edge directions : Cliques have cliquewidth 2, tournaments have unbounded clique-width.
(3) Deciding “Clique-width < 3” is a polynomial problem. (Habib et al.)
The complexity (polynomial or NP-complete) of “Clique-width = 4” is unknown.
It is NP-complete to decide for given k and G if cwd(G) < k. (Fellows et al.)
There exists a cubic approximation algorithm that for given k and G
answers (correctly) :
either that cwd(G) >k,
or produces a clique-width algebraic term using 224k labels. (Oum)
This yields Fixed Parameter Tractable algorithms for many hard problems.
Example : Cliques have clique-width 2.
Kn is defined by tn where tn+1 = Relabb,a(Add-edga,b(tn  b))
Another example : Cographs
Cographs are generated by  and  defined by :
G H
= Relabb,a (Add-edga,b (G  Relaba,b(H))
= G  H with “all edges” between G and H.
5. Algorithmic applications
Fixed parameter tractability results
Theorem (B.C.) : A) For graphs of clique-width  k ,
each monadic second-order property, (ex. 3-colorability),
each monadic second-order optimization function, (ex. distance),
each monadic second-order counting function, (ex. # of paths)
is evaluable :
in linear time on graphs given by a term over VR-operations,
in time O(n3) otherwise (by S. Oum, 2005).
B) All this is possible in linear time on graphs of tree-width  k, for each fixed
k (by Bodlaender, 1996).
Inductive computations
Example : Series-parallel graphs, defined as graphs with sources 1 and 2,
generated from e = 1
and the operations // (parallel-composition) and
series-composition defined from other operations by :
G  H = Forget3(Ren2,3 (G) // Ren1,3 (H))
Example :
 
Inductive proofs :
G, H connected implies : G//H and G  H are connected, (induction)
e is connected (basis) :
All series-parallel graphs are connected.
3) It is not true that :
G and H planar implies : G//H is planar (K5 = H//e).
A stronger property for induction :
G has a planar embedding with the sources in the same “face”
All series-parallel graphs are planar.
Inductive computation : Test for 2-colorability
1) Not all series-parallel graphs are 2-colorable (see K3)
2) G, H 2-colorable does not imply that G//H is 2-colorable (because K3=P3//e).
3) One can check 2-colorability with 2 auxiliary properties :
Same(G) = G is 2-colorable with sources of the same color,
Diff(G) = G is 2-colorable with sources of different colors
by using rules : Diff(e) =True ; Same(e) = False
Same(G//H)  Same(G) Same(H)
Diff(G//H)  Diff(G)  Diff(H)
Same(GH)  (Same(G) Same (H))  (Diff(G) Diff(H))
Diff(GH)  (Same(G) Diff(H))  (Diff(G)  Same(H))
We can compute for every SP-term t, by induction on the structure of t the pair of
Boolean values (Same(Val(t)) , Diff(Val(t)) ).
We get the answer for G = Val(t) (the graph that is the value of t ) regarding 2colorability.
Important facts :
1) The existence of properties forming an inductive set (w.r.t. operations of
F) is equivalent to recognizability in the considered F-algebra.
2) The simultaneous computation of m inductive properties can be
implemented by a "tree" automaton with 2m states working on terms t. This
computation takes time O(t).
3) An inductive set of properties can be constructed (at least theoretically)
from every monadic-second order formula.
4) This result extends to the computation of values (integers) defined by
monadic-second order formulas.
Recognizability and inductive properties
Definition : A finite set P of properties on an F-algebra M is F-inductive
if for
every p P and f F, there exists a (known) Boolean formula B such that :
p(fM(a,b) ) = B[…,q(a),…,q(b),….,qP]
for all a and b in M.
(where q(a),…, q(b)  {True, False}) .
Proposition : A subset L of M is recognizable iff it is the set of elements that
satisfy a property belonging to a finite inductive set of properties P.
Proof : Let L = h-1(C) for a homomorphism h : M  A , A a finite F-algebra and C a
subset of A (domain of A).
For each a in A, let â be the property : â(m)= True  h(m) = a. Let p be such
that p(m) = True
 h(m) C m  L.
Properties {p, â / aA} form an F-inductive set.
If P is an inductive set of k properties, one can define an F-algebra structure on the
set Bk of k-tuples of Booleans, such that the mapping
h : m  the k-tuple of Booleans is a homomorphism.
Inductive properties and automata on terms
P is an inductive set of k properties, one can define a deterministic
automaton on terms of T(F) with set of states the k-tuples of Booleans, that
computes in a bottom-up way, for each term t, the truth values :
p(Val(s)) for all p  P and all subterms s of t.
Membership of an element m of M in a recognizable set L can be tested
by such an automaton on any term t in T(F) defining m.
Application to graphs
Immediate but depends on two things :
Parsing algorithms building terms from the given graphs :
1) results by Bodlaender for constructing tree-decompositions (in linear time),
whence terms representing them,
2) results by Oum for constructing (non-optimal) terms for graphs of cliquewidth at most k. (Cubic time algorithm).
properties : Monadic second-order
6. Monadic Second-Order (MS) Logic
A logical language which specifies inductive properties and functions
= First-order logic on power-set structures
= First-order logic extended with (quantified) variables
denoting subsets of the domains.
MS properties : transitive closure, properties of paths, connectivity,
planarity (via Kuratowski, uses connectivity), k-colorability.
Examples of formulas for G = < VG , edgG(.,.) >, undirected
Non connectivity :
X ( x  X  y X & u,v (u  X  edg(u,v)  v  X)
2-colorability (i.e. G is bipartite) :
X ( u,v (u  X  edg(u,v)  v  X)  u,v (u  X  edg(u,v)  v  X) )
Edge set quantifications
Incidence graph of G undirected, Inc(G) = < VG  EG, incG(.,.) >.
incG(v,e)  v is a vertex of edge e.
Monadic second-order (MS2) formulas written with inc can use quantifications
on sets of edges.
Existence of Hamiltonian circuit is expressible by an MS2 formula, but not by an MS
Theorem : MS2 formulas are no more powerful than MS formulas :
for graphs of degree < d, or of tree-width < k,
or for planar graphs, or graphs without some fixed H as a minor,
or graphs of average degree < k (uniformly k-sparse).
Definition : A set L of words, of trees, of graphs or relational structures is
Monadic Second-Order (MS) definable iff
L = {S /
S   } for an MS formula 
Theorem : (1) A language (set of words or finite terms ) is
recognizable  it is MS definable
(2) A set of finite graphs is VR- or VR+-recognizable 
 it is MS definable
(3) A set of finite graphs is HR-recognizable 
 it is MS2 definable
(1) Doner, Thatcher, Wright, see W. Thomas, Handbook formal languages, vol.3.
(2) (3) There are two possible proofs .
Basic facts for (2) :
Let F consist of  and unary quantifier-free definable operations f.
For every MS formula
of quantifier-height k, we have
(a) for every f , one can construct a formula f#() such that :
f(S)    S  f#()
(b) (Hanf, Fefermann and Vaught, Shelah) one can construct formulas
such that :
ST    for some i, S  i  T   i
where f#(), 1,…,n, 1,…,n have quantifier-height < k.
(c) Up to equivalence, there are finitely many formulas of quantifier-height < k forming a set
One builds an automaton with states the subsets of
k : the MS-theories of quantifier-height < k of
the graphs defined by the subterms of the term to be processed.
7. Monadic second-order transductions
STR(): the set of finite
-relational structures (or finite directed ranked -hypergraphs).
MS transductions are multivalued mappings :
: STR()  STR()
S T = (S)
where T is :
a) defined by MS formulas
b) inside the structure: S  S  ...  S
(fixed number of disjoint "marked" copies of S)
c) in terms of "parameters", subsets X1, …,Xp of the domain of S.
Proposition : The composition of two MS transductions is an MS
Remark : For each tuple of parameters X1, …,Xp satisfying an MS property, T is
uniquely defined.  is multivalued by the different choices of parameters.
Examples : (G,{x}) the connected component containing x.
(G,X,Y) the minor of G resulting from contraction of edges in X and deletion of
edges and vertices in Y.
Example of an MS transduction (without parameters) : The square mapping
on words: u  uu
u =
aac, we have
In (S)
   
   
   
       
we redefine Suc (i.e.,  ) as follows :
Suc(x,y) :  p1 (x) & p1 (y) & Suc(x,y) v p2 (x) & p2 (y) & Suc(x,y)
v p1 (x) & p2 (y) & "x has no successor" & "y has no predecessor"
We also remove the "marker" predicates p1, p2.
The fundamental property of MS transductions :
S 
 
Every MS formula  has an effectively computable backwards translation (),
an MS formula, such that :
S = () iff (S) = 
The verification of  in the object structure (S) reduces to the verification of
() in the given structure S.
Intuition : S contain all necessary information to describe (S) ; the MS properties of (S)
are expressible by MS formulas in S
Consequence : If L  STR() has a decidable MS satisfiability problem, so has its image
under an MS transduction.
Another look at MS2
versus MS formulas
Theorem : The mapping G  Inc(G) is an MS-transduction
on each of the following classes of simple graphs :
degree < d,
tree-width < k,
planar graphs,
graphs without some fixed H as a minor,
graphs of average degree < k (uniformly k-sparse) .
Other results
1) A set of graphs is VR -equational iff it is the image of (all) binary trees under an MS
transduction. VR-equational sets are stable under MS-transductions.
A set of graphs has bounded clique-width iff it is the image of a set of binary
trees under an MS transduction.
2) A set of graphs is HR-equational iff it is the image of (all) binary trees under an
MS2 transduction.
HR-equational sets are stable under MS2-transductions.
A set of graphs has bounded tree-width iff it is the image of a set of binary trees
under an MS2 transduction.
3) A set of hypergraphs is QF-equational iff it is the image of (all) binary trees under
an MS-transduction.
QF-equational sets are stable under MS-transductions.
4) (A.Blumensath, B.C., 2004) : QF-recognizable sets are preserved under inverse MS
Relationships between algebraic and logical notions
notions characterizations characterizations
union, Rec
equation systems MS-trans(Trees)
Boolean opns
MS-def REC
Signatures for graphs and hypergraphs :
HR :
VR :
VR+ :
QF :
graphs and hypergraphs with “sources”
graphs with vertex labels (“ports”)
VR with quantifier-free operations (ex. edge complement)
hypergraphs, i.e., relational structures (disjoint union  and
quantifier-free definable unary operations)
8. Links between MS logic and combinatorics: Seese’s Theorem
and Conjecture
Theorem (Seese 1991): If a set of graphs has a decidable MS2 satisfiability
problem, it has bounded tree-width.
Conjecture (Seese 1991): If a set of graphs has a decidable MS satisfiability
problem, it is the image of a set of trees under an MS transduction,
equivalently, has bounded clique-width.
Theorem (B.C., S. Oum 2004): If a set of graphs has a decidable C2MS
satisfiability problem, it has bounded clique-width.
MS = (Basic) MS logic without edge quantifications, MS2 = MS logic with edge quantifications
C2MS = MS logic with even cardinality set predicates. A set C has a decidable L satisfiability
problem if one can decide for every formula in L whether it is satisfied by some graph in C
Proof of Seese’s Theorem :
A) If a set of graphs C has unbounded tree-width, the set of its minors includes
all k x k-grids (Robertson, Seymour)
B) If a set of graphs contains all kxk-grids, its MS2 satisfiability problem is
C) If C has decidable MS2 satisfiability problem, so has Minors(C),
because C
Minors(C) is an MS2 transduction.
Hence, if C has unbounded tree-width and a decidable MS2 satisfiability
problem, we have a contradiction for the decidability of the MS2 satisfiability problem
of Minors(C).
Proof of Courcelle-Oum’s Theorem :
D) Equivalence between the cases of all (directed and undirected) graphs and
bipartite undirected graphs.
A’) If a set of bipartite graphs C has unbounded clique-width, the set of its vertexminors contains all “Sk“ graphs
C’) If C has decidable C2MS satisfiability problem, so has Vertex-Minors(C),
because C
Vertex-Minors(C) is a C2MS transduction.
E) An MS transduction transforms Sk into the kxk-grid.
Hence A' + B + C' + E gives the result for bipartite undirected graphs. Result with D.
Definitions and facts
Local complementation of G at vertex v
G * v = G with edge complementation of G[nG(v)],
the subgraph induced by the neighbours of v
Local equivalence (  loc ) = transitive closure of local complementation
(at all vertices)
Vertex-minor relation :
H <VM G :  H is an induced subgraph of some G’  loc G.
Proposition (Courcelle and Oum 2007) : The mapping that associates with G its locally
equivalent graphs is a C2MS transduction.
Why is the even cardinality set predicate necessary ?
Consider G * X for X Y :
u is linked to v in G * X
 Card(X) is even
Definition of Sk : bipartite : A = {1,…,(k+1)(k-1)} , B = {1,…,k(k-1)} for j  A, i  B :
edg(i,j )  i  j  i+k-1
From Sk to Gridkxk by an MS transduction
(folded) Grid3x4
1) One can define the orderings of A and B :
x, y are consecutive  Card(nG(x)  nG(y)) = 2
2) One can identify the edges from i
 B to i  A, and
from i  B to i+k-1  A (thick edges on the left drawing)
3) One can create edges (e.g. from 1  A to 2  A, from 2  A to 3  A etc…and
similarly for B, and from 1  B to 4  A, etc…) and delete others (from 4  B to 6 A
etc…), and vertices like 7,8 in A, to get a grid containing Gridkxk
Corollary : If a set of directed acyclic graphs having Hamiltonian directed paths has a
decidable MS satisfiability problem, then :
it has bounded clique-width,
it is the image of a set of trees under an MS transduction.
Proof : Since on these graphs a linear order is MS definable, MS and C2MS are
The previously known techniques for similar results (in particular for line graphs or
interval graphs, B.C. 2004) do not work in this case.
9. A few open questions
Question 1 (A. Blumensath, P. Weil, B.C.): Which operations, quantifier-free definable or
not, yield extensions of VR, HR, QF that are equivalent ?
Question 2 : Under which operations, quantifier-free definable or not, are REC(VR) and
REC(HR) closed ?
The case of REC(HR) is considered in B.C.: (HR-)Recognizable sets of graphs,
equivalent definitions and closure properties, 1994. It is not hard to see that REC(VR)
is closed under  (disjoint union) but not under the operations Add-edga,b.
Answer : unary operations with inverse defined by an MS-transduction.
Question 3 : Is it true that the decidability of the MS (and not of the C2MS) satisfiability
problem for a set of graphs implies bounded clique-width, as conjectured by D. Seese?
What about sets of hypergraphs or relational structures?

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