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Alexandre Madeira
Mário Benevides (Eds.)
LNCS 10669
Dynamic Logic
New Trends and Applications
First International Workshop, DALI 2017
Brasilia, Brazil, September 23–24, 2017
Proceedings
123
Lecture Notes in Computer Science 10669
Commenced Publication in 1973
Founding and Former Series Editors:
Gerhard Goos, Juris Hartmanis, and Jan van Leeuwen
Editorial Board
David Hutchison
Lancaster University, Lancaster, UK
Takeo Kanade
Carnegie Mellon University, Pittsburgh, PA, USA
Josef Kittler
University of Surrey, Guildford, UK
Jon M. Kleinberg
Cornell University, Ithaca, NY, USA
Friedemann Mattern
ETH Zurich, Zurich, Switzerland
John C. Mitchell
Stanford University, Stanford, CA, USA
Moni Naor
Weizmann Institute of Science, Rehovot, Israel
C. Pandu Rangan
Indian Institute of Technology, Madras, India
Bernhard Steffen
TU Dortmund University, Dortmund, Germany
Demetri Terzopoulos
University of California, Los Angeles, CA, USA
Doug Tygar
University of California, Berkeley, CA, USA
Gerhard Weikum
Max Planck Institute for Informatics, Saarbrücken, Germany
More information about this series at http://www.springer.com/series/7407
Alexandre Madeira Mário Benevides (Eds.)
•
Dynamic Logic
New Trends and Applications
First International Workshop, DALI 2017
Brasilia, Brazil, September 23–24, 2017
Proceedings
123
Editors
Alexandre Madeira Mário Benevides
INESC TEC UFRJ
University of Minho Rio de Janeiro
Braga Brazil
Portugal
and
CIDMA
University Aveiro
Aveiro
Portugal
Building on the pioneer intuitions of Floyd–Hoare logic, dynamic logic was introduced
in the 1970s by Vaughan Pratt as a suitable logic to reason about, and verify, classic
imperative programs. Since then, the original intuitions grew to an entire family of
logics, which became increasingly popular for assertional reasoning about a wide range
of computational systems. Simultaneously, their object (i.e., the very notion of a
program) evolved in unexpected ways. This leads to dynamic logics tailored to specific
programming paradigms and extended to new computing domains, including proba-
bilistic, continuous, and quantum computation. Both its theoretical relevance and
practical potential make dynamic logic a topic of interest in a number of scientific
venues, from wide-scope software engineering conferences to modal logic-specific
events. However, as far as we know, to date, no specific event was exclusively dedi-
cated to it. This workshop emerged from this discussion, during the kick-off meeting
of the project DaLí - Dynamic Logics for Cyber-Physical Systems, whose editors are
participating as consultant and IR, respectively.
This volume contains the proceedings of the event that was held in the beautiful city
of Brasilia during September 23–24, 2017, co-located with Tableaux, ITP, and FroCoS
2017.
We received 24 submissions, of which, after a careful revision process, with at least
three revisions per work, 12 papers were accepted as regular papers and are published
in this volume. Beyond these contributions, the workshop also included the following
short papers:
– Fabricio Chalub, Alexandre Rademaker, Edward Hermann Haeusler and Christiano
Braga: “Fixing the Proof of completeness of ALC Sequent Calculus”
– Diana Costa and Édi Duarte: “Checkers Game in Deontic Logic”
– Daniel Figueiredo and Manuel A. Martins: “Bisimulations for Reactive Frames”
– Konstantinos Gkikas and Alexandru Baltag: “Stable Beliefs and Conditional
Probability Spaces”
– Leandro Gomes: “Contract-Based Design for Software Verification”
– Luiz Carlos Pereira: “Constructive Fragments of Classical Modal Logic and the
Ecumenical Perspective”
– Carlos Tavares: “Toward a Quantum-Probabilistic Dynamic Logic”
From this list, four papers were originally submitted as short contribution and the
other three were regular submissions invited to be converted into a short format. An
informal volume with the extended abstracts (3–5 pages) of these contributions was
provided in the conference.
We also had the invited talks of Alexandru Baltag, “Logic Goes Viral: Dynamic
Modalities for Social Networks,” and of Edward Hermann Haeusler, “Propositional
Dynamic Logic with Petri Net Programs: A Discussion and a Logical System”; the
latter is a joint work with Bruno Lopes and Mario Benevides. Finally, we had the
VI Preface
Conference Chairs
Mario Benevides UFRJ, Brazil
Alexandre Madeira University of Minho, Portugal
Program Committee
Carlos Areces University of Cordoba, Argentina
Phillippe Balbiani University of Toulouse, France
Alexandru Baltag Uva, The Netherlands
Luís S. Barbosa University of Minho, Portugal
Johan van Benthem University of Stanford, USA and University of Tsinghua,
China
Patrick Blackburn University of Roskilde, Denmark
Fredrik Dahlqvist UCL, UK
Stéphane Demri CNRS, France
Hans van Ditmarsch LORIA, Nancy, France
Francicleber M. Ferreira UFC, Brazil
Valentin Goranko University of Stockholm, Sweden
Edward H. Hauesler PUC-Rio, Brazil
Rolf Hennicker LMU, Munich, Germany
Andreas Herzig Toulouse, France
Dexter Kozen Cornell, USA
Clemens Kupke University of Strathclyde, UK
Bruno Lopes Vieira UFF, Brazil
Paulo Mateus IST, Portugal
Manuel A. Martins University of Aveiro, Portugal
Carlos Olarte UF RN, Brazil
José N. Oliveira University of Minho, Portugal
André Platzer CMU, USA
Eugénio Rocha University of Aveiro, Portugal
Valéria de Paiva NC, USA
Regivan Santiago UFRN, Brazil
Luis Menasche UFRJ, Brazil
Schechter
Tinko Tinchev Sofia University, Bulgaria
Petrucio Viana UFF, Brazil
Sheila Veloso UFRJ, Brazil
Yde Venema ILLC, The Netherlands
Renata Wassermann USP, Brazil
VIII Organization
Additional Reviewers
Brandon Bohrer
Diana Costa
Daniel Figueiredo
Isaque Lima
Vitor Machado
Johannes Marti
António Pereira
Elaine Pimentel
Sponsoring
This workshop was promoted by DaLí, a research project financed by the ERDF
European Regional Development Fund through the Operational Programme for
Competitiveness and Internationalization, COMPETE 2020 Program, and by National
Funds through the Portuguese funding agency, FCT, Fundação para a Ciência e a
Tecnologia, within reference POCI-01-0145-FEDER-016692.
FCT, Fundação para a Ciência e Tecnologia
CIDMA, Center for Research and Development in Mathematics and Applications
High Assurance Laboratory INESC TEC
Contents
1 Introduction
Modal logics [12,14] were originally conceived as logics of necessary and possible
truths. They are now viewed, more broadly, as logics that explore a wide range
of modalities, or modes of truth: epistemic (“it is known that”), doxastic (“it is
believed that”), deontic (“it ought to be the case that”), or temporal (“it has
been the case that”), among others. From a model theoretic perspective, the
field evolved into a discipline that deals with languages interpreted on various
kinds of relational structures or graphs. Nowadays, modal logics are actively used
in areas as diverse as software verification, artificial intelligence, semantics and
pragmatics of natural language, law, philosophy, etc.
From an abstract point of view, modal logics can be seen as formal languages
to navigate and explore properties of a given relational structure. But if we
want to describe and reason about dynamic aspects of a given situation, e.g.,
how the relations between a set of elements evolve through time or through
the application of certain operations, the use of modal logics (or actually, any
kind of logic with classical semantics) becomes less clear. We can always resort
c Springer International Publishing AG 2018
A. Madeira and M. Benevides (Eds.): DALI 2017, LNCS 10669, pp. 1–16, 2018.
https://doi.org/10.1007/978-3-319-73579-5_1
2 C. Areces et al.
to modeling the whole space of possible evolutions as a graph, but this soon
becomes unwieldy. It would be more elegant to use truly dynamic modal logics
with operators that can mimic the changes the structure will undergo.
There exist several dynamic modal operators that fit in this approach. A
clear example are the dynamic operators introduced in dynamic epistemic logics
(see, e.g. [22]). Less obvious examples are given by hybrid logics [8,13] equipped
with the down arrow operator ↓ which is used to ‘rebind’ names for states to the
current point of evaluation, and memory logics [19], a kind of restricted form
of hybrid logics that come equipped with a memory and operators to store and
retrieve states from it. Finally, a classical example which can arguably be taken
as the origin of the studies of logics in this approach is Sabotage Logic introduced
by van Benthem in [21], which provides an operator that deletes individual edges
in the model.
Generalizing this last logic, we study operators that do various kinds of
change to the accessibility relation of a model: deleting, adding, and swapping
edges, both locally (near the state of evaluation) and globally (anywhere). We
call these operators relation-changing. In [2], the operators are introduced, and
it is shown that the model checking problem is PSPACE-complete for the basic
modal logic enriched with any of these operators. In this article, we consider
the satisfiability problem of these logics. Previous results on this topic are the
undecidability of (multimodal) global sabotage logic, via encoding of the Post
Correspondence Problem [16] the undecidability of local swap logic with a sin-
gle relation, by reduction from memory logic [4]; and non-terminating tableau
methods for all six logics [3]. Here we present undecidability proofs for all six
logics using reductions from memory logic.
The undecidability results can be surprising, considering for instance that
dynamic epistemic logics are decidable [11,17,22]. However, other very expres-
sive dynamic operators are undecidable, such as the hybrid logic with the ↓
operator [8]. As we mentioned before, ↓ binds states of the model to some par-
ticular names. We will show in this article that relation-changing operators can
take advantage of adding, deleting or swapping around edges, to perform some
sort of binding in the model, turning them undecidable.
Contributions
– We sketch the undecidability proof for the memory logic ML( r , k ), by adapt-
ing the undecidability argument introduced in [18] for the description logic
ALCself.
– We introduce undecidability proofs for the satisfiability problem of six relation-
changing modal logics via satisfiability of memory logic. In this way, we com-
plete the picture of the computational aspects of the family of languages
defined in this framework.
– Our proofs improve previous ones for local swap [4] and global sabotage [16],
by exploiting undecidability of memory logics. This allows for shorter proofs
and avoid redundant encodings of the tiling problem.
The article is organized as follows. In Sect. 2 we introduce the syntax and
semantics of relation-changing modal logics. In Sect. 3 we introduce the memory
Undecidability of Relation-Changing Modal Logics 3
where p ∈ PROP, ∈ {sb, br, sw, gsb, gbr, gsw} and ϕ, ψ ∈ FORM.
Other operators are defined as usual. In particular, ϕ is defined as ¬¬ϕ.
Let ML (the basic modal logic) be the logic without the {sb, br, sw, gsb,
gbr, gsw} operators, and ML() the extension of ML allowing also , for
∈ {sb, br, sw, gsb, gbr, gsw}.
(sabotaging) M− − −
S = W, RS , V , with RS = R\S, S ⊆ R.
(bridging) MS = W, RS , V , with RS = R ∪ S, S ⊆ (W ×W )\R.
+ + +
Intuitively, M−
S is obtained from M by deleting the edges in S, and similarly
M+
S adds the edges in S to the accessibility relation, and M∗S adds the edges in
S as inverses of edges previously in the accessibility relation.
Let w be a state in M, the pair (M, w) is called a pointed model (we will
usually drop parentheses). In the rest of this article, we will use wv as a shorthand
for {(w, v)} or (w, v); context will always disambiguate the intended use.
4 C. Areces et al.
M, w |= p iff w ∈ V (p)
M, w |= ¬ϕ iff M, w |= ϕ
M, w |= ϕ ∧ ψ iff M, w |= ϕ and M, w |= ψ
M, w |= ♦ϕ iff for some v ∈ W s.t. (w, v) ∈ R, M, v |= ϕ
M, w |= sbϕ iff for some v ∈ W s.t. (w, v) ∈ R, M− wv , v |= ϕ
M, w |= brϕ iff for some v ∈ W s.t. (w, v) ∈ R, M+ wv , v |= ϕ
M, w |= swϕ iff for some v ∈ W s.t. (w, v) ∈ R, M∗vw , v |= ϕ
M, w |= gsbϕ iff for some v, u ∈ W, s.t. (v, u) ∈ R, M−vu , w |= ϕ
M, w |= gbrϕ iff for some v, u ∈ W, s.t. (v, u) ∈ R, M+vu , w |= ϕ
M, w |= gswϕ iff for some v, u ∈ W, s.t. (v, u) ∈ R, M∗uv , w |= ϕ.
FORM ::= p | k | ¬ϕ | ϕ ∧ ψ | ♦ϕ | r ϕ,
The remaining cases coincide with the semantics of ML, and do not involve the
memory.
An ML( r , k )-formula ϕ is satisfiable if there are a model M = W, R, V, ∅
and w ∈ W such that M, w |= ϕ. The empty initial memory ensures that no state
of the model satisfies the unary predicate k unless a formula r ψ has previously
been evaluated there.
(¬ k ∧ r ψ) ∨ ( k ∧ ψ)
Sabotage Logic
Local Sabotage. In the translation to local sabotage logic, the Structsb sub-
formula should ensure that every state of the model can be memorized using the
expressivity of sb. This operator changes the point of evaluation after deleting
an edge. To compensate for this, the Structsb formula guarantees that every
state has an edge that is deleted when the state is memorized, and a second edge
back to the original state to ensure that evaluation can continue at the correct
state. We use a spy point s to ensure this structure. The idea is illustrated in
the following image.
ϕ
...
s
We need to ensure that every satisfiable formula of ML( r , k ) is translated
into a satisfiable formula (and vice-versa, if the translated formula is satisfiable,
then the original formula is satisfiable, too). The image above shows an intended
model for the translated formula τsb (ϕ). Intuitively, bold edges and arrows
correspond to the model of ϕ. The complete translation is given in Definition 8.
Here we discuss in detail how it works.
Structsb adds a spy state with symmetric edges between itself and all other
states. In particular, (1) in Definition 8 ensures that the evaluation state satisfies
s and that it is irreflexive, and (2) guarantees that its immediate successors
reach a state where s holds. Formulas (3) and (4) ensure that this state is the
Undecidability of Relation-Changing Modal Logics 7
original s state. They work together as follows: (3) makes ♦s true in any s-
state reachable in two steps, and by deleting the traversed edges we avoid a cycle
of size two between this s-state and an immediate successor of the evaluation
state, distinguishing the original s-state from any other s-state reachable in two
steps. (4) then traverses one edge, deletes the next one, and reaches a state where
s implies ♦¬s. This contradicts (3), unless we have arrived in the original s
state. Formulas (5), (6) and (7) mimic (2), (3) and (4), but for edges which are
removed twice. Observe that (6) now avoids a cycle of size three between any
other s-state reachable in two steps and an immediate successor of the evaluation
state. Finally, (8) and (9) ensure that the evaluation state is indeed a spy state,
i.e., that it is linked to every other state of the input model.
Trsb starts by placing the translation ( ) of ϕ in a successor of the evaluation
state. Boolean cases are obvious. For the diamond case, ♦ψ is satisfied if there
is a successor v where ψ holds, but we must ensure that v is not the spy state.
For ( r ψ) , we do a round-trip of sabotaging from the current state to the spy
state. Note that after reaching the spy state an edge does come back to the same
state where it came from, since the only accessible state where ¬♦s holds is the
one we are memorizing. For ( k ) , we check whether there is an edge pointing to
some s-state.
Lemma 1. Let ϕ be an ML( r , k )-formula in PNF that does not contain the
propositional symbol s. Then, ϕ is satisfiable iff τsb (ϕ) is satisfiable.
8 C. Areces et al.
...
ϕ
Lemma 2. Let ϕ be an ML( r , k )-formula in PNF that does not contain the
propositional symbol s. Then, ϕ is satisfiable iff τgsb (ϕ) is satisfiable.
Bridge Logic
Local Bridge. For local bridge logic, we use a spy state that is initially discon-
nected from the input model. When some state should be memorized, the spy
state gets connected (in both directions) to it. This construction is quite special
since we do not have pre-built gadgets in the input model, as they get built on
demand.
10 C. Areces et al.
Let us first show the following result, that enables us to force the evaluation
state to be the only one in the model to satisfy s:
Proof. First, w obviously satisfies s and does not have any successor. Now, we
have M, w |= [br](s→[br]¬s). In particular this means that M+ ww , w |= s→[br]¬s,
hence M+ ww , w |= [br]¬s. Since in M +
ww , the state w is only connected to itself,
this means that for all v = w, we have M+ ww,wv , v |= ¬s, this also means that
M, v |= s for all v = w.
For Bridge Logics, Structbr adds to the input model a spy state in which s
holds. By Lemma 3, (1) in Definition 10 ensures that the evaluation state has
no successor and is the only state in the model where s holds. And (2) ensures
that there are no edges from ¬s-states (anywhere in the model) to the spy state.
The idea is illustrated in the following image, where t is a propositional symbol
used in Trbr (ϕ) and dotted lines represent edges created with the br operator.
t ...
ϕ
Trbr (ϕ) first creates two edges until a ¬s-state, where the translation of ϕ
holds. For Trbr ( r ) we do a round-trip of bridging from the current state to
the spy state. Note that the second part of this round-trip has to be from the
spy state to the remembered state, since it is the only way to satisfy br(♦s).
Also note that this would not work if the s state was directly connected to the
input model; this is why we use the intermediate t-state. For Trbr ( k ) we check
whether there is an edge to a state where s holds.
Proposition 4. Let W, R, V a model such that there is a unique state s where
s holds, there is no state x ∈ W such that (x, s) ∈ R, and there is a component
Undecidability of Relation-Changing Modal Logics 11
Lemma 4. Let ϕ be an ML( r , k )-formula in PNF that does not contain the
propositional symbols s and t. Then, ϕ is satisfiable iff τbr (ϕ) is satisfiable.
Global Bridge. The global bridge operator is able to add edges in the model.
This is why, to mark some state, we use this operator to add an edge to some
s-state. Then, we enforce that the initial model does not have any reachable
s-state.
Here Structgbr (ϕ) ensures that no state of the input model has s-successors.
Storing a state in the memory is simulated by creating an edge to an s-state, and
checking whether the current state of evaluation is in the memory is simulated
by checking the presence of an s-successor. Observe that we could have either
one state where s holds or (possibly) different s-states for each state of the input
model.
s
...
ϕ
12 C. Areces et al.
Definition 11. Define τgbr (ϕ) = Structgbr (ϕ) ∧ Trgbr (ϕ), where:
Structgbr (ϕ) = i ¬s
0≤i≤md(ϕ)+1
Swap Logic
Local Swap. We introduce a new version of the translation given in [4] that uses
only one propositional symbol. The idea is that we have each state pointing
to some states called switch states, and memorizing a state is represented by
swapping such edges. Then, no edge pointing to a switch
means that the state
has been memorized. We use the notation (n) ϕ for i ϕ.
1≤i≤n
In this case Structsw adds “switch states”, which are in one-to-one corre-
spondence with the states of the input model, together with a spy state. By (2)
in Definition 12, each ¬s-state at one, two and three steps from the evaluation
state, has a unique dead-end successor where s holds (switch state). By (3) and
(4), switch states (corresponding to states at distance 1, 2 and 3) can be reached
from the evaluation state by a unique path. (5) makes the evaluation state a spy
state. All these conjuncts together ensure that switch states are independent one
from another. The idea is illustrated in the following image.
s
s s
...
ϕ
s
Definition 12. Define τsw = Structsw ∧ Trsw (ϕ), where:
Structsw =
s ∧ ¬s (1)
∧ (3) (¬s → U niq) (2)
∧ [sw](s → (s → ⊥) (3)
∧ [sw](s → (s → ⊥) (4)
∧ [sw][sw](¬s→sw(s ∧ ♦((¬s)→♦♦(s ∧ ♦¬♦s)))) (5)
For Trsw ( r ϕ) we traverse and swap the edge between the current state
and its switch state, and come back to the same state. For Trsw ( k ), we check
whether the current state has not an edge to its switch state.
Lemma 6. Let ϕ be an ML( r , k )-formula in PNF that does not contain the
propositional symbol s. Then, ϕ is satisfiable iff τsw (ϕ) is satisfiable.
Proof. (⇐) From a pointed model W, R, V , w of τsw (ϕ) we can extract a
pointed model W , R , V , ∅, w satisfying ϕ following the same definition as in
the proof of Lemma 1.
For all ψ sub-formula of ϕ, v ∈ W , S ⊆ W , T = {(v , v) | v ∈ S ∧ (v, v ) ∈
R ∧ v ∈ V (s)} and RS = (R\T −1 ) ∪ T , we will prove that W , R , V , S, v |= ψ
if, and only if, W, RS , V , v |= (ψ) .
We do it by structural induction on ψ. We prove the ¬ k ∧ r χ case. Suppose
W , R , V , S, v |= ¬ k ∧ r χ. Then by definition, v ∈ / S and W , R , V , S ∪
{v}, v |= χ, and by Proposition 5, we have (v, v ) ∈ RS for a unique v ∈ V (s).
Global Swap. The global swap operator is able to change the direction of some
edge in the model. In particular, we are interested in the ability to swap, for
some state, an incoming edge (undetectable for the basic modal logic) into an
outgoing edge. This is why this translation is similar to the one of global bridge
logic. Initially, the model does not have any reachable state where s holds. As
14 C. Areces et al.
for global sabotage and global bridge, there may be many states where s holds
in the model with edges to states of the input model. The idea is illustrated in
the following image, where only one s state is shown.
s
...
ϕ
Definition 13. Define τgsw (ϕ) = Structgsw (ϕ) ∧ Trgsw (ϕ), where:
Structgsw (ϕ) = i ¬s
0≤i≤md(ϕ)+1
Lemma 7. Let ϕ be an ML( r , k )-formula in PNF that does not contain the
propositional symbol s. Then, ϕ is satisfiable iff τgsw (ϕ) is satisfiable.
5 Conclusions
We exploited the similarities between memory logic and relation-changing logics
to obtain simple and non-redundant undecidability proofs. We first presented an
undecidability result for memory logics in the monomodal case, by adapting the
proof introduced in [18] for ALCself. Then, we presented translations from the
satisfiability problem of monomodal memory logics to all six relation-changing
modal logics. Both results combined show undecidability of the satisfiability
problem for relation-changing modal logics in a very simple way. These results
complete the picture of the computational behaviour of relation-changing logics,
given that we already know that model checking for them is PSPACE-complete
[2,4,5,15].
This high complexity of the logics is a consequence of the degree of liberty
we give to the operators. By replacing arbitrary modifications with conditional
modifications (i.e., according to a pre- and a post-condition) it is possible to
decrease the complexity and get decidable logics (e.g., as in [9,10]).
Undecidability of Relation-Changing Modal Logics 15
A related problem is the one of finite satisfiability. Indeed, for many applica-
tions of dynamic epistemic logic, we are only interested in looking for finite mod-
els. Finite satisfiability is known to be undecidable for multimodal global
sabotage logic [20], and decidable for monomodal local sabotage and local swap
logics [6]. It remains to see the status of this problem for all remaining cases.
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Axiomatization and Computability of a Variant
of Iteration-Free P DL with Fork
1 Introduction
Propositional dynamic logic (P DL) is an applied non-classical logic designed
for reasoning about the behaviour of programs [10]. The definition of its syntax
is based on the idea of associating with each program α of some programming
language the modal operator [α], formulas of the form [α]φ being read “every
execution of the program α from the present state leads to a state bearing
the formula φ”. Completeness and decidability results for the standard version
of P DL in which programs are built up from program variables and tests by
means of the operations of composition, union and iteration are given in [15,16].
A number of interesting variants have been obtained by extending or restricting
the syntax or the semantics of P DL in different ways [7,9,14,18].
Some of these variants extend the ordinary semantics of P DL by considering
sets W of states structured by means of a function from the set of all pairs of
states into the set of all states [5,11–13]: the state x is the result of applying the
function to the states y, z iff the information concerning x can be separated in
a first part concerning y and a second part concerning z. The binary function
considered in [5,11] has its origin in the addition of an extra binary operation of
fork denoted ∇ in relation algebras: in [5, Sect. 2], whenever x and y are related
via R and z and t are related via S, states in x z and states in y t are related
via R∇S whereas in [11, Chap. 1], whenever x and y are related via R and x
and z are related via S, x and states in y z are related via R∇S.
This addition of fork in relation algebras gives rise to a variant of P DL
which includes the program operation of fork denoted Δ. In this variant, for all
programs α and β, one can use the modal operator [αΔβ], formulas of the form
[αΔβ]φ being read “every execution in parallel of the programs α and β from the
c Springer International Publishing AG 2018
A. Madeira and M. Benevides (Eds.): DALI 2017, LNCS 10669, pp. 17–34, 2018.
https://doi.org/10.1007/978-3-319-73579-5_2
18 P. Balbiani and J. Boudou
present state leads to a state bearing the formula φ”. The binary operation of
fork ∇ considered in Benevides et al. [5, Sect. 2] gives rise to P RSP DL, a variant
of P DL with fork whose axiomatization is still open. We devote this paper to
the axiomatization and the computability of P DLΔ 0 , a variant of iteration-free
P DL with fork whose semantics is based on the interpretation of the binary
operation of fork ∇ considered in Frias [11, Chap. 1].
The difficulty in axiomatizing or deciding P RSP DL or P DLΔ 0 originates in
the fact that the program operations of fork considered above are not modally
definable in the ordinary language of P DL. We overcome this difficulty by means
of tools and techniques developed in [1,3,4]. Our results are based on the fol-
lowing: although fork is not modally definable, it becomes definable in a modal
language strengthened by the introduction of propositional quantifiers. Instead
of using axioms to define the program operation of fork in the language of P DL
enlarged with propositional quantifiers, we add an unorthodox rule of proof that
makes the canonical model standard for the program operation of fork and we
use large programs for the proof of the Truth Lemma.
We will first present the syntax (Sect. 2) and the semantics (Sect. 3) of P DLΔ 0
and continue with results concerning the expressivity of P DLΔ 0 (Sect. 4), the
axiomatization/completeness of P DLΔ 0 (Sects. 5 and 6) and the decidability of
P DLΔ 0 (Sect. 7). We assume the reader is at home with tools and techniques in
modal logic and dynamic logic. For more on this, see [6,15]. The proofs of our
results can be found in [2].
2 Syntax
Definition 2 (Abbreviations). The modal constructs for formulas ··, (·¯ ◦·),
(·¯·) and (·¯·) are defined as follows: αφ ::= ¬[α]¬φ; (φ¯
◦ψ) ::= ¬(¬φ ◦ ¬ψ);
(φ¯ψ) ::= ¬(¬φ ¬ψ); (φ¯ψ) ::= ¬(¬φ ¬ψ). Moreover, for all formulas φ, let
φ0 ::= ¬φ and φ1 ::= φ.
Axiomatization and Computability of a Variant of Iteration-Free PDL 19
It is well worth noting that programs and formulas are finite strings of sym-
bols coming from a countable alphabet. It follows that there are countably many
programs and countably many formulas. The construct ·; · comes from the class
of algebras of binary relations [19]: the program α; β firstly executes α and sec-
ondly executes β. As for the construct ·Δ·, it comes from the class of proper fork
algebras [11, Chap. 1]: the program αΔβ performs a kind of parallel execution
of α and β. The construct [·]· comes from the language of P DL [10,15]: the for-
mula [α]φ says that “every execution of α from the present state leads to a state
bearing the information φ”. As for the constructs · ◦ ·, · · and · ·, they come
from the language of conjugated arrow logic [8,17]: the formula φ ◦ ψ says that
“the present state is a combination of states bearing the information φ and ψ”,
the formula φ ψ says that “the present state can be combined to its left with a
state bearing the information φ giving us a state bearing the information ψ” and
the formula φ ψ says that “the present state can be combined to its right with
a state bearing the information ψ giving us a state bearing the information φ”.
Example 1. The formula [aΔb](p ◦ q) says that “the parallel execution of a and
b from the present state always leads to a state resulting from the combination
of states bearing the information p and q”.
Definition 3 (Test insertion). Let f be the function from the set of all pro-
grams into itself inductively defined as follows:
– f (a) = a;
– f (α; β) = f (α); ?; f (β);
– f (αΔβ) = (f (α); ?)Δ(f (β); ?);
– f (φ?) = φ?.
CHAPTER IV.
Discovery of the Pacific Ocean.—Plans of Columbus.—Avarice of
the Spaniards.—Balboa.—Weighing the gold.—The young
Indian’s speech.—Indian mode of fighting.—Balboa ascends
the mountain.—First view of the Pacific.
Columbus had first seen land in the New World on the 12th of
October, 1492. Six years after he surveyed the coast of the
American continent by Paria and Cumana. Territory was the grand
object with the noble mind of Columbus; he wished to colonize this
great country by the settling of Europeans, and thus introduce
Christianity and civilization among the Red Men. But the adventurers
that followed him sought gold as their only object, and employed the
sword as the only means of converting the natives.
The Spaniards who first landed on the continent, saw before them
a magnificent country, vast forests, mighty rivers, long ranges of
mountains—a dominion wide enough for the widest ambition of
conquest, or the richest enjoyment of life; but no treasure. Still their
avarice was kept in a perpetual fever by the Indian stories of gold in
profusion farther to the west, and their fancy was excited by tales of
a sea beyond, which they said stretched to the extremities of the
globe.
The first European who set his eye on the Pacific Ocean, was
Vasco Thenez De Balboa. His family was of the order of Spanish
gentry. He was a man of great enterprise, personal strength, and of a
daring courage. He had been disappointed in his expectations of
obtaining wealth at Hayti, where he had settled, and an expedition
sailing to Darien, he accompanied it. A colony was already
established on the eastern side of the isthmus of Darien; but the
savages in the vicinity had been found so warlike, that the settlers
did not venture to explore the interior.
Indian rumors of the golden country continued to inflame the
Spaniards. They heard of one king Dabaibe, who was said to be
living in a city filled with treasure, and who worshipped an idol of
solid gold. Balboa put himself at the head of his countrymen, and
marched to conquer the rich city. But they had first to conquer the
surrounding caciques, who would not permit the Spaniards to pass
through their territories. At length, Balboa formed an alliance with
Comogre, a mountain chieftain, who had three thousand warriors.
The son of Comogre brought a present to the Spanish troops of
sixty slaves and four thousand pieces of gold. In distributing the gold,
some difficulty occurred, as is usually the case where people are all
selfish; the quarrel grew furious, and swords were drawn. The young
Indian looked on, first with astonishment, then with scorn. Advancing
to the scales in which they were weighing the gold, he threw them on
the ground, exclaiming—“Is it for this trifle that you Spaniards
quarrel? If you care for gold, go seek it where it grows. I can show
you a land where you may gather it by handfuls.”
This speech brought all the Spaniards around him, and he
proceeded to detail his knowledge. “A cacique, very rich in gold,”
said he, “lives to the south, six suns off.” He pointed in that direction.
“There,” said he, “you will find the sea. But there you will find ships
as large as your own, with sails and oars. The men of these lands
are so rich, that their common eating and drinking vessels are of
gold.” This was to the Spaniards their first knowledge of Peru.
Balboa determined to search for this rich country. He collected a
hundred and ninety Spanish soldiers, a thousand friendly Indians,
and some bloodhounds, and began his march into the wilderness.
The Indian tribes were instantly roused. The Spaniards had scarcely
reached the foot of the Sierra, when they found the warriors, headed
by their caciques, drawn up in a little army.
The Indians, like the ancient Greeks, first defied the enemy, by
loud reproaches and expressions of scorn. They then commenced
the engagement. Torecha, their king, stood forth in the front of his
people, clothed in a regal mantle, and gave the word of attack. The
Indians rushed on with shouts; but the Spanish crossbows and
muskets were terrible weapons to their naked courage. The Indians
were met by a shower of arrows and balls, which threw them into
confusion. They were terrified, also, at the noise of the guns. They
thought the Spaniards fought with thunder and lightning. Still, the
Indians did not fly till their heroic king and six hundred of their
warriors were left dead on the spot. Over their bleeding bodies,
Balboa marched to the plunder of their city.
Balboa, with his army, now commenced the ascent of the
mountains. It took them twenty days. After toiling through forests,
and climbing mountains that seemed inaccessible, his Indian guide
pointed out to him, among the misty summits of the hills that lay
before him, the one from which the Pacific was visible. Balboa
determined to have the glory of looking upon it first. He commanded
his troops to halt at the foot of the hill. He ascended alone, with his
sword drawn, and having reached the summit, cast his eyes around.
The Pacific Ocean was spread out before him!
Balboa had invaded the Indian country in search of gold, and
murdered the natives to obtain it; but at that time such conduct was
not considered very wicked. The Indians were looked upon with
horror, because they were savages, and Balboa believed himself a
good Christian because he was a Catholic. He fell on his knees, and,
weeping, offered his thanksgiving to Heaven, for the bounty that had
suffered him to see this glorious sight. He doubtless thought God
was well pleased with him.
His troops had watched his ascent of the mountain, with the
eagerness of men who felt their fates bound up in his success. When
they saw his gestures of delight and wonder, followed by his falling
on his knees and prayer, they became incapable of all restraint. They
rushed up the hill like wild deer. But when they saw the matchless
prospect around them, they, too, shared the spirit of their leader;
they fell on their knees and offered up their thanksgiving to God. Yet
at the same time they doubtless contemplated plundering and
destroying the Indians. They had not learned to do to others as they
would have others do to them.
Lion Hunting.
Most people are more disposed to run away from lions than to
run after them, unless indeed they are safely locked up in cages. But
only think of going to hunt lions in the wilderness! Yet such things are
done in Africa, where lions are frequently met with.
In the southern part of that country is a tribe of negroes called the
Bechuana. The men of this tribe are accustomed to carry a long staff
with a bunch of ostrich feathers tied at one end, which is used to
shade themselves from the sun. It is in fact a kind of parasol, but
whether it is designed to save their complexion, I cannot say. It
seems, at any rate, that the ladies do not use it. But beside serving
as a parasol, this feathered staff has another and important use. As I
have said, these people sometimes go in pursuit of the lion, and
when a party of hunters meet one, they go near to him, and as he
springs on one of them, the hunter quickly plants the handle of the
staff in the ground and retreats. The fierce lion leaps upon the staff
and rends the ostrich feathers in pieces. While he is thus engaged,
the other hunters come suddenly upon him from behind, and
despatch him with their daggers.
CHAPTER IX.
Completion of my education.—Manly sports.—An accident.—The
bed of pain.—Recovery from sickness.—A new companion.