Book ChapterDOI

# Decidable extensions of Church's problem

07 Sep 2009-pp 424-439
TL;DR: A parameterized version of the Church synthesis problem, in this extended version a formula B and a finite-state operator F might contain as a parameter a unary predicate P, such that the Church problem with the parameter P is decidable.
Abstract: For a two-variable formula B(X,Y) of Monadic Logic of Order (MLO) the Church Synthesis Problem concerns the existence and construction of a finite-state operator Y=F(X) such that B(X, F(X)) is universally valid over Nat. Buchi and Landweber (1969) proved that the Church synthesis problem is decidable. We investigate a parameterized version of the Church synthesis problem. In this extended version a formula B and a finite-state operator F might contain as a parameter a unary predicate P. A large class of predicates P is exhibited such that the Church problem with the parameter P is decidable. Our proofs use Composition Method and game theoretical techniques.

### 1 Introduction

• Two fundamental results of classical automata theory are decidability of the monadic second-order logic of order (MLO) over ω = (N, <) and computability of the Church synthesis problem.
• In [11, 14] the authors provided necessary and sufficient conditions for the decidability of monadic (second-order) theory of expansions of the linear order of the naturals ω by unary predicates.
• The authors can compute for the winning player in Gωϕ a finite- state winning strategy.
• The authors main results show that the synthesis problem for each predicate P ∈ ER is decidable.

### 2 Preliminaries and Background

• The authors use N for the set of natural numbers and ω for the first infinite ordinal.
• The authors use the expressions “chain” and “linear order” interchangeably.
• Well formed formulas of the monadic logic MLO are obtained from atomic formulas using Boolean connectives ¬,∨,∧,→ and the first-order quantifiers ∃t and ∀t, and the secondorder quantifiers ∃X and ∀X.

### 2.2 Elements of the composition method

• The authors proofs make use of the technique known as the composition method developed by Feferman-Vaught and Shelah [8, 17].
• So, there are finitely many ≡n-equivalence classes of l-structures.
• The authors will use only special cases of this definition in which the index chain I and the summand chains.
• The function which maps the pairs of characteristic formulas to their sum is a recursive function.
• The authors often use the following well-known lemmas: Lemma 2.8.

### 3 Game types

• The authors define games on game types and show that these game are reducible to the McNaughton games.
• But first the authors introduce a terminology, define finite-memory strategies and fix some notational conventions.
• Whenever M is clear from the context the authors write typem(ρ) for typem(M ⌢ρ).
• Notational Conventions 1. In Hintikka’s Lemma the authors considered formulas with the free variables among X1, . . . ,Xl. Definition 3.1.
• The authors consider the following ω-game Game(F,G). Game(F,G): The following proposition plays an important role in their proofs: Proposition 3.4.

### 4 Winning strategies over classes of finite chains

• In this subsection the authors will consider the games over expansions of finite chains.
• The only difference is that these games are of finite length.
• The games over an l-chains with m elements have m rounds.
• ⊓⊔ Proposition 4.6 is crucial for the design of their algorithm, due the decidability of (4).

### 5 Algorithm

• For every MLO formula ϕ(X1,X2, P ), first construct a set of the characteristic formulas G such that ϕ is equivalent to their disjunction and then use the following algorithm.
• From U and {V l}∞l=0 the authors can compute the desirable Out.
• Usually, this property fails for ER predicates; however, the sequence of the distances modulo n behaves periodically.
• The authors algorithm is more subtle than the above sketch for Pex and relies on this periodicity.
• The rest of the game Player I will play according to his finite-memory strategy st3(F,G2) computed in the Step 3 Clearly, the described strategy is a finite-memory strategy.

### 6 Further Results and Open Questions

• The authors proved that the finite-memory synthesis problem is decidable for the expansions of ω by the predicates from ER.
• The game GMϕ (h1, h2) with look-ahead h1 for Player I and look-ahead h2 for Player II is defined as follows.
• The proof of the next proposition is similar to the proof of Theorem 1.4.
• There is an algorithm that for every MLO formula ϕ(X1,X2, Z) decides whether Player I has a finite-memory winning strategy in GMϕ (h1, h2), and if so, constructs such a strategy.
• It is plausible that in their proofs the compositional methods can be hidden and a presentation can be given based on automata theoretic concepts.

Did you find this useful? Give us your feedback

Content maybe subject to copyright    Report

Decidable Extensions of Church’s Problem
Alexander Rabinovich
The Blavatnik School of Computer Science, Tel Aviv University, Israel
rabinoa@post.tau.ac.il
Abstract. For a two-variable formula B(X,Y) of Monadic Logic of Or-
der (MLO) the Church Synthesis Problem concerns the existence and
construction of a ﬁnite-state operator Y=F(X) such that B(X,F(X)) is
universally valid over Nat.
B ¨uchi and Landweber (1969) proved that the Church synthesis problem
is decidable.
We investigate a parameterized version of the Church synthesis problem.
In this extended version a formula B and a ﬁnite-state operator F might
contain as a parameter a unary predicate P.
A large class of predicates P is exhibited such that the Church problem
with the parameter P is decidable.
Our proofs use Composition Method and game theoretical techniques.
1 Introduction
Two fundamental results of classical automata theory are decidability of the
monadic second-order logic of order (MLO) over ω = (N, <) and computability
of the Church synthesis problem. These results have provided the underlying
mathematical framework for the development of formalisms for the description of
interactive systems and their desired properties, the algorithmic veriﬁcation and
the automatic synthesis of correct implementations from logical speciﬁcations,
and advanced algorithmic techniques that are now embodied in industrial tools
for veriﬁcation and validation.
Decidable Expansions of ω B¨uchi [1] proved that the monadic theory of
ω = (N, <) is decidable. Even before the decidability of the monadic theory
of ω has been proved, it was shown that the expansions of ω by “interesting”
of (N, <, +) and the monadic theory of (N, <, λx.2 × x) are undecidable [15, 20].
Therefore, most eﬀorts to ﬁnd decidable expansions of ω deal with expansions
Elgot and Rabin [5] found many interesting predicates P for which MLO over
(N, <, P ) is decidable. Among these predicates are the set of factorial numbers
{n! | n N}, the sets of k-th powers {n
k
| n N} and the sets {k
n
| n N} (for
k N ).
The Elgot and Rabin method has been generalized and sharpened over the
years and their results were extended to a variety of unary predicates (see e.g.,
[18, 16, 3]). In [11, 14] we provided necessary and suﬃcient conditions for the
decidability of monadic (second-order) theory of expansions of the linear order
of the naturals ω by unary predicates.

Church’s Problem What is known as the “Church synthesis problem” was
ﬁrst posed by A. Church in [4] for the case of (ω, <). The Church problem is
much more complicated than the decidability problem for M LO. Church uses
the language of automata theory. It was McNaughton (see [9]) who ﬁrst observed
that the Church problem can be equivalently phrased in game-theoretic language
and in recent years many authors took up the generalizations of such games for
various applications of the algorithmic theory of inﬁnite games (see e.g., [6, 10]).
McNaughton considered games over ω. We consider such games over expansions
of ω by unary predicates.
Let M = (N, <, P ) be the expansion of ω by a unary predicate P . Let
ϕ(X
1
, X
2
, Z) be a formula, where X
1
, X
2
and Z are set (monadic predicate)
variables. The McNaughton game G
M
ϕ
is deﬁned as follows.
1. The game is played by two players, called Player I and Player II.
2. A pla y of the game has ω rounds.
3. At round i N: ﬁrst, Player I chooses ρ
X
1
(i) {0, 1}; then, Player II chooses
ρ
X
2
(i) {0, 1}. Both players can observe whether i P .
4. By the end of the play two predicates ρ
X
1
, ρ
X
2
N have been constructed
1
5. Then, Player I wins the play if M |= ϕ(ρ
X
1
, ρ
X
2
, P ); otherwise, Player II
wins the play.
What we want to know is: Does either one of the players have a winning strategy
in G
M
ϕ
? If so, which one? That is, can Player I choose his moves so that, whatever
way Player II responds we have ϕ(ρ
X
1
, ρ
X
2
, P )? Or can Player II respond to
Player I’s moves in a way that ensures the opposite?
At round i, Player I has access only to ρ
X
1
(0) . . . ρ
X
1
(i1), ρ
X
2
(0) . . . ρ
X
2
(i
1) and P (0) . . . P (i).
Hence, a strategy of Player I can be deﬁned as a function which assigns to
any ﬁnite sequence
(ρ
X
1
(0), ρ
X
2
(0), P (0)) . . . (ρ
X
1
(i 1), ρ
X
2
(i 1), P (i 1)) (, , P (i))
a value in {0, 1} which is taken to be ρ
X
1
(i).
At round i, Player II has access only to ρ
X
1
(0) . . . ρ
X
1
(i), ρ
X
2
(0) . . . ρ
X
2
(i1)
and P (0) . . . P (i).
Hence, a strategy of Player II can be deﬁned as a function which assigns to
any ﬁnite sequence
(ρ
X
1
(0), ρ
X
2
(0), P (0)) . . . (ρ
X
1
(i 1), ρ
X
2
(i 1), P (i 1)) (ρ
X
1
(i), , P (i))
a value in {0, 1} which is taken to be ρ
X
2
(i).
Since strategies are functions from ﬁnite strings (over a ﬁnite alphabet) to
{0, 1} we can classify them according to their complexity. The recursive strate-
gies, the ﬁnite-memory strategies, i.e., the strategies computable by ﬁnite-state
transducers are deﬁned in a natural way (see Sect. 3).
1
We identify monadic predicates with their characteristic functions.

We investigate the following parameterized version of the Church synthesis
problem.
Synthesis Problems for M = (N, <, P ), where P N
Input: an MLO formula ϕ(X
1
, X
2
, Z).
Task: Check whether Player I has a ﬁnite-memory winning strategy in G
M
ϕ
and if there is such a strategy - construct it.
To simplify notations, games and the synthesis problem were previously de-
ﬁned for formulas with three free variables X
1
, X
2
and Z. It is easy to generalize
all deﬁnitions and results to formulas ψ(X
1
, . . . , X
m
, Y
1
, . . . Y
n
, Z
1
, . . . , Z
l
) with
many variables. In this generalization at r ound β, Player I chooses values for
X
1
(β), . . . , X
m
(β), then Player II replies by choosing the values to Y
1
(β), . . . , Y
n
(β)
and the structure M provides the interpretation for Z
1
, . . . Z
l
. Note that, strictly
speaking, the input to the synthesis problem is not only a formula, but a for-
mula plus a partition of its free-variables to Player I’s variables and Player II’s
variables and parameter’s variables.
In [2], uchi and Landweber prove the computability of the synthesis problem
in ω = (N, <) (no parameters).
Theorem 1.1 (B¨uchi-Landweber, 1969). Let ϕ(
¯
X,
¯
Y ) be a formula, where
¯
X and
¯
Y are disjoint lists of variables. Then:
Determinacy: One o f the players has a winning strategy in the game G
ω
ϕ
.
Decidability: It is decidable which of the players has a winning strategy.
Finite-state strategy: The player who has a winning strategy, als o has a ﬁnite-
state winning strategy.
Synthesis algorithm: We can compute for the winning player in G
ω
ϕ
a ﬁnite-
state winning strategy.
The determinacy part of the theorem follows from the topological arguments.
In particular for every expansion M of ω by unary predicates, the game G
M
ϕ
is
determinate.
Let M be an expansion of ω by unary predicates. We proved in [12], that
there is an algorithm which for every MLO formula ϕ decides who wins G
M
ϕ
if
and only if the monadic theory of M is decidable. Moreover, we proved that if the
monadic theory of M is decidable, then the player who has a winning strategy in
G
M
ϕ
has a recursive MLO-deﬁnable winning strategy which is computable from
ϕ.
The ﬁnite-state strategy part of Theorem 1.1 fails for decidable expansions
of ω. For example, let Fac = {n! | n N} be the set of factorial numbers. The
fac
:= (N, <, Fac) is decidable by [5]. Let ϕ(X
1
, X
2
, Z) be
a formula which speciﬁes that t X
1
iﬀ t + 1 Z (hence for the game G
M
fac
ϕ
the
moves of Player II are irrelevant). It is easy to see that Player I has a winning
strategy in G
M
fac
ϕ
, yet Player I has no ﬁnite-state winning strategy in this game.
The results of this paper imply that the synthesis problem for (N, <, Fac) is
decidable.

Main Result Our main result describes a large class of predicates P such that
the synthesis problem for (N, <, P ) is decidable.
An ω-sequence a
i
is said to be ultimately periodic with lag l and period d if
a
i
= a
i+d
for i > l.
Deﬁnition 1.2. Let
¯
k = (k
1
< k
2
< . . . k
i
< . . . ) be an increasing ω-sequence
of integers.
1.
¯
k is sparse if for each d there is n such that k
i+1
k
i
> d for each i > n.
¯
k is eﬀectively sparse if there is an algorithm that for each d computes n
such that k
i+1
k
i
> d for each i > n.
2.
¯
k is ultimately reducible if for every m > 1 the sequence k
i
mod m is ulti-
mately periodic.
¯
k is eﬀectively ultimately reducible if there is an algorithm
that for each m computes a lag and a period of k
i
mod m.
Deﬁnition 1.3. Let ER be the class of increasing recursive ω-sequences of in-
tegers which are eﬀectively sparse and eﬀectively ultimately reducible.
Let P N be a predicate. We denote by Enum(P ) the sequence (k
1
, k
2
. . . k
i
. . . )
which enumerates the elements of P in the increasing order. Often we do not
distinguish between P and Enum(P ), In particular we say that a predicate is
ER predicate if Enum(P ) is in ER. The class ER contains many interesting
predicates. It contains the set Fact={n! | n N} of factorial numbers, the sets
{k
n
| n N}, the sets {n
k
| n N}. It has nice closure properties, e.g. if
¯
k and
¯
l are in ER then {k
i
+ l
i
| i N}, {k
i
× l
i
| i N}, and {k
l
i
i
| i N} are in ER.
In [18], Siefkes introduced ER predicates and generalized Elgot-Rabin con-
traction method to prove that for every ER predicate P the monadic theory of
M = (N, <, P ) is decidable. Our main results show that the synthesis problem
for each predicate P ER is decidable.
Theorem 1.4 (Main). Let P be an ER predicate and let M = (N, <, P ).
There is an algorithm that for every MLO formula ϕ(X
1
, X
2
, Z) decides whether
Player I has a ﬁnite-memory winning strategy in G
M
ϕ
, and if so constructs such
a st rategy.
Our algorithm is based on game theoretical techniques and the composition
method developed by Feferman-Vaught, Shelah and others.
Organization of the paper The article is organized as follows. The next sec-
and summarizes elements of the composition method. In Section 3, we introduce
game-types, deﬁne games on game types and show that these game are reducible
to the McNaughton games. Section 4 consider games over ﬁnite chains. Suﬃcient
conditions are provided for existence of a ﬁnite state strategies which uniformly
wins over a class of ﬁnite chains.
Section 5 describes an algorithm for the synthesis problem over the expan-
sions of ω by ER predicates, and proves the soundness of the algorithm, i.e., if
the algorithm outputs a strategy for G
M
ϕ
, then it is a ﬁnite state strategy which
wins ϕ over M. The proof of completeness appears in the full version of this
paper [13]. Further results and open questions are discussed in Sect. 6.

2 Preliminaries and Background
We use i, j, n, k, l, m, p, q for natural numbers. We use N for the set of natural
numbers and ω for the ﬁrst inﬁnite ordinal. We use the expressions chain and
linear order interchangeably. A chain with m elements will be denoted by m.
We use P(A) for the set of subsets of A.
2.1 The Monadic Logic of Order (MLO )
Syntax The syntax of the monadic second-order logic of order - MLO has in
its vocabulary individual (ﬁrst order ) variables t
1
, t
2
variables X
1
, X
2
. . . and one binary relation < (the order).
Atomic formulas are of the form X(t) and t
1
< t
2
. Well formed formulas
of the monadic logic MLO are obtained from atomic formulas using Boolean
connectives ¬, , , and the ﬁrst-order quantiﬁers t and t, and the second-
order quantiﬁers X and X. The quantiﬁer depth of a formula ϕ is denoted by
qd(ϕ).
We use upper case letters X, Y , Z,... to denote second-order variables; with
an overline,
¯
X,
¯
Y , etc., to denote ﬁnite tuples of variables.
Semantics A structure is a tuple M := (A, <
M
,
¯
P
M
) where: A is a non-empty
set, <
M
is a binary relation on A, and
¯
P
M
:=
P
M
1
, . . . , P
M
l
is a ﬁnite tuple
of subsets of A.
If
¯
P
M
is a tuple of l sets, we call M an l-structure. If <
M
linearly orders A,
we call M an l-chain. When the speciﬁc l is unimportant, we simply say that
M is a labeled chain.
Suppose M is an l-structure and ϕ a formula with free-variables among
X
1
, . . . , X
l
. We deﬁne the relation M |= ϕ (read: M satisﬁes ϕ) as usual, un-
derstanding that the s econd-order quantiﬁers range over subsets of A.
Let M be an l-structure. The monadic theory of M, MTh(M), is the set of
all formulas with free-variables among X
1
, . . . , X
l
satisﬁed by M.
From now on, we omit the superscript in <
M
and
¯
P
M
’. We often write
(A, <) |= ϕ(
¯
P ) meaning (A, <,
¯
P ) |= ϕ.
For a chain M = (A, <,
¯
P ) and a subset I of A, we denote by M
I the
subchain of M over the set I.
2.2 Elements of the composition method
Our proofs make use of the technique known as the composition method de-
veloped by Feferman-Vaught and Shelah [8, 17]. To ﬁx notations and to aid the
reader unfamiliar with this technique, we brieﬂy review the deﬁnitions and re-
sults that we require. A more detailed presentation can be found in [19] or [7].
Let n, l N. We denote by Form
n
l
the set of MLO formulas with free variables
among X
1
, . . . , X
l
and of quantiﬁer depth n.
Deﬁnition 2.1. Let n, l N and let M, N be l-structures. The n-theory of M
is Th
n
(M) := {ϕ Form
n
l
| M |= ϕ}. If Th
n
(M) = Th
n
(N ), we say that M
and N are n-equivalent and write M
n
N .

##### Citations
More filters
01 Jan 2011
TL;DR: A k-ExpTime algorithm is presented to compute global positional winning strategies for parity games which are played on the configuration graph of a level-k higher-order pushdown system and the solution of games in the sense of Gale and Stewart where the winning condition is specified by an MSO-formula φ(P ) with a parameter P ⊆ N.
Abstract: Higher-order pushdown systems extend the idea of pushdown systems by using a “higher-order stack” (which is a nested stack). More precisely on level 1 this is a standard stack, on level 2 it is a stack of stacks, and so on. We study the higher-order pushdown systems in the context of infinite regular games. In the first part, we present a k-ExpTime algorithm to compute global positional winning strategies for parity games which are played on the configuration graph of a level-k higher-order pushdown system. To represent those winning strategies in a finite way we use a notion of regularity for sets of higher-order stacks that relies on certain (“symmetric”) operations to build higher-order stacks. The construction of the strategies is based on automata theoretic techniques and uses the fact that the higher-order stacks constructed by symmetric operations can be arranged uniquely in a tree structure. In the second part, we study the solution of games in the sense of Gale and Stewart where the winning condition is specified by an MSO-formula φ(P ) with a parameter P ⊆ N. This corresponds to a three player game where the i-th round between the two original players is extended by the choice of the bit 1 or 0 depending on whether i ∈ P or not. We consider the case that the parameter can be constructed by some deterministic machine, a “parameter generator”. We solve the parametrized regular games for parameters P given by two kinds of such generators, namely: higher-order pushdown automata and collapsible pushdown automata. In the third part, we study higher-order pushdown systems and higherorder counter systems (where the stack alphabet contains only one symbol), by comparing the language classes accepted by corresponding automata. For example, we show that level-k pushdown languages are level-(k+1) counter languages.

3 citations

Journal ArticleDOI
TL;DR: A parameterized version of the Church synthesis problem, in this extended version a formula B and a finite-state operator F might contain as a parameter a unary predicate P, such that the Church problem with the parameter P is decidable.
Abstract: For a two-variable formula B(X,Y) of Monadic Logic of Order (MLO) the Church synthesis problem concerns the existence and construction of a finite-state operator Y=F(X) such that B(X,F(X)) is universally valid over Nat. Buchi and Landweber (1969) proved that the Church synthesis problem is decidable. We investigate a parameterized version of the Church synthesis problem. In this extended version a formula B and a finite-state operator F might contain as a parameter a unary predicate P. A large class of predicates P is exhibited such that the Church problem with the parameter P is decidable. Our proofs use composition method and game theoretical techniques.

3 citations

### Cites background from "Decidable extensions of Church's pr..."

• ...An extended abstract of this paper was published in [13]....

[...]

Proceedings ArticleDOI
14 Jul 2022
TL;DR: The decidability of checking whether the ﬁrst player has a winning strategy in the realizability game for a given CLTL formula is studied, and it is proved thatDecidability is maintained for single-sided games, even if the formulas are allowed to be prompt-CLTL formulas.
Abstract: Constraint linear-time temporal logic (CLTL) is an extension of LTL that is interpreted on sequences of valuations of variables over an infinite domain. The atomic formulas are interpreted as constraints on the valuations. The atomic formulas can constrain valuations over a range of positions along a sequence, with the range being bounded by a parameter depending on the formula. The satisfiability and model checking problems for CLTL have been studied by Demri and D’Souza. We consider the realizability problem for CLTL. The set of variables is partitioned into two parts, with each part controlled by a player. Players take turns to choose valuations for their variables, generating a sequence of valuations. The winning condition is specified by a CLTL formula – the first player wins if the sequence of valuations satisfies the specified formula. We study the decidability of checking whether the first player has a winning strategy in the realizability game for a given CLTL formula. We prove that it is decidable in the case where the domain satisfies the completion property, a property introduced by Balbiani and Condotta in the context of satisfiability. We prove that it is undecidable over ( Z , <, =), the domain of integers with order and equality. We prove that over ( Z , <, =), it is decidable if the atomic constraints in the formula can only constrain the current valuations of variables belonging to the second player, but there are no such restrictions for the variables belonging to the first player. We call this single-sided games.
##### References
More filters
Book ChapterDOI
01 Jan 1990
TL;DR: The interpreted formalism of SC as mentioned in this paper is a fraction of the restricted second order theory of natural numbers, or of the first-order theory of real numbers, and it is easy to see that SC is equivalent to the first order theory [Re, +, Pw, Nn], whereby Re, + are the sets of non-negative reals, integral powers of 2, and natural numbers.
Abstract: Let SC be the interpreted formalism which makes use of individual variables t, x, y, z,... ranging over natural numbers, monadic predicate variables q( ), r( ), s( ), i( ),... ranging over arbitrary sets of natural numbers, the individual symbol 0 standing for zero, the function symbol ′ denoting the successor function, propositional connectives, and quantifiers for both types of variables. Thus SC is a fraction of the restricted second order theory of natural numbers, or of the first order theory of real numbers. In fact, if predicates on natural numbers are interpreted as binary expansions of real numbers, it is easy to see that SC is equivalent to the first order theory of [Re, +, Pw, Nn], whereby Re, Pw, Nn are, respectively, the sets of non-negative reals, integral powers of 2, and natural numbers.

1,325 citations

### "Decidable extensions of Church's pr..." refers background in this paper

• ...Theorem 1.1 (Büchi-Landweber, 1969)....

[...]

• ...Decidable Expansions of ω Büchi [1] proved that the monadic theory of ω = (N, <) is decidable....

[...]

• ...Decidable Expansions of ω Büchi [1] proved that the monadic theory of ω = (N, ) is decidable....

[...]

• ...In [2], Büchi and Landweber prove the computability of the synthesis problem in ω = (N, ) (no parameters)....

[...]

Book ChapterDOI
21 Aug 2000
TL;DR: In the ambient logic of classical second order propositional calculus, the specification problem for a family of excluded middle like tautologies is solved and these are shown to be realized by sequential simulations of specific communication schemes for which they provide a safe typing mechanism.
Abstract: In the ambient logic of classical second order propositional calculus, we solve the specification problem for a family of excluded middle like tautologies. These are shown to be realized by sequential simulations of specific communication schemes for which they provide a safe typing mechanism.

1,119 citations

Book
01 Jan 2002
TL;DR: The 19 chapters presented in this multi-author monograph give a consolidated overview of the research results achieved in the theory of automata, logics, and infinite games during the past 10 years.
Abstract: A central aim and ever-lasting dream of computer science is to put the development of hardware and software systems on a mathematical basis which is both firm and practical. Such a scientific foundation is needed especially for the construction of reactive programs, like communication protocols or control systems. For the construction and analysis of reactive systems an elegant and powerful theory has been developed based on automata theory, logical systems for the specification of nonterminating behavior, and infinite two-person games. The 19 chapters presented in this multi-author monograph give a consolidated overview of the research results achieved in the theory of automata, logics, and infinite games during the past 10 years. Special emphasis is placed on coherent style, complete coverage of all relevant topics, motivation, examples, justification of constructions, and exercises.

646 citations

Book ChapterDOI
TL;DR: In this article, the authors present an algorithm which decides whether or not a condition X, Y stated in sequential calculus admits a finite automata solution, and produces one if it exists.
Abstract: Our main purpose is to present an algorithm which decides whether or not a condition 𝕮(X, Y) stated in sequential calculus admits a finite automata solution, and produces one if it exists. This solves a problem stated in [4] and contains, as a very special case, the answer to Case 4 left open in [6]. In an equally appealing form the result can be restated in the terminology of [7], [10], [15]: Every ω-game definable in sequential calculus is determined. Moreover the player who has a winning strategy, in fact, has a winning finite-state strategy, that is one which can effectively be played in a strong sense. The main proof, that of the central Theorem 1, will be presented at the end. We begin with a discussion of its consequences.

597 citations

### "Decidable extensions of Church's pr..." refers background in this paper

• ...In [2], Büchi and Landweber prove the computability of the synthesis problem in ω = (N, <) (no parameters)....

[...]

BookDOI
01 Jan 2002

523 citations